2005-04-17 02:20:36 +04:00
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#ifndef _LINUX_MMZONE_H
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#define _LINUX_MMZONE_H
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#ifndef __ASSEMBLY__
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2008-04-28 13:12:54 +04:00
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#ifndef __GENERATING_BOUNDS_H
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2005-04-17 02:20:36 +04:00
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#include <linux/spinlock.h>
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#include <linux/list.h>
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#include <linux/wait.h>
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2007-10-17 10:25:54 +04:00
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#include <linux/bitops.h>
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2005-04-17 02:20:36 +04:00
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#include <linux/cache.h>
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#include <linux/threads.h>
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#include <linux/numa.h>
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#include <linux/init.h>
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2005-10-30 04:16:53 +03:00
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#include <linux/seqlock.h>
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2006-03-27 13:15:57 +04:00
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#include <linux/nodemask.h>
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Add a bitmap that is used to track flags affecting a block of pages
Here is the latest revision of the anti-fragmentation patches. Of particular
note in this version is special treatment of high-order atomic allocations.
Care is taken to group them together and avoid grouping pages of other types
near them. Artifical tests imply that it works. I'm trying to get the
hardware together that would allow setting up of a "real" test. If anyone
already has a setup and test that can trigger the atomic-allocation problem,
I'd appreciate a test of these patches and a report. The second major change
is that these patches will apply cleanly with patches that implement
anti-fragmentation through zones.
kernbench shows effectively no performance difference varying between -0.2%
and +2% on a variety of test machines. Success rates for huge page allocation
are dramatically increased. For example, on a ppc64 machine, the vanilla
kernel was only able to allocate 1% of memory as a hugepage and this was due
to a single hugepage reserved as min_free_kbytes. With these patches applied,
17% was allocatable as superpages. With reclaim-related fixes from Andy
Whitcroft, it was 40% and further reclaim-related improvements should increase
this further.
Changelog Since V28
o Group high-order atomic allocations together
o It is no longer required to set min_free_kbytes to 10% of memory. A value
of 16384 in most cases will be sufficient
o Now applied with zone-based anti-fragmentation
o Fix incorrect VM_BUG_ON within buffered_rmqueue()
o Reorder the stack so later patches do not back out work from earlier patches
o Fix bug were journal pages were being treated as movable
o Bias placement of non-movable pages to lower PFNs
o More agressive clustering of reclaimable pages in reactions to workloads
like updatedb that flood the size of inode caches
Changelog Since V27
o Renamed anti-fragmentation to Page Clustering. Anti-fragmentation was giving
the mistaken impression that it was the 100% solution for high order
allocations. Instead, it greatly increases the chances high-order
allocations will succeed and lays the foundation for defragmentation and
memory hot-remove to work properly
o Redefine page groupings based on ability to migrate or reclaim instead of
basing on reclaimability alone
o Get rid of spurious inits
o Per-cpu lists are no longer split up per-type. Instead the per-cpu list is
searched for a page of the appropriate type
o Added more explanation commentary
o Fix up bug in pageblock code where bitmap was used before being initalised
Changelog Since V26
o Fix double init of lists in setup_pageset
Changelog Since V25
o Fix loop order of for_each_rclmtype_order so that order of loop matches args
o gfpflags_to_rclmtype uses gfp_t instead of unsigned long
o Rename get_pageblock_type() to get_page_rclmtype()
o Fix alignment problem in move_freepages()
o Add mechanism for assigning flags to blocks of pages instead of page->flags
o On fallback, do not examine the preferred list of free pages a second time
The purpose of these patches is to reduce external fragmentation by grouping
pages of related types together. When pages are migrated (or reclaimed under
memory pressure), large contiguous pages will be freed.
This patch works by categorising allocations by their ability to migrate;
Movable - The pages may be moved with the page migration mechanism. These are
generally userspace pages.
Reclaimable - These are allocations for some kernel caches that are
reclaimable or allocations that are known to be very short-lived.
Unmovable - These are pages that are allocated by the kernel that
are not trivially reclaimed. For example, the memory allocated for a
loaded module would be in this category. By default, allocations are
considered to be of this type
HighAtomic - These are high-order allocations belonging to callers that
cannot sleep or perform any IO. In practice, this is restricted to
jumbo frame allocation for network receive. It is assumed that the
allocations are short-lived
Instead of having one MAX_ORDER-sized array of free lists in struct free_area,
there is one for each type of reclaimability. Once a 2^MAX_ORDER block of
pages is split for a type of allocation, it is added to the free-lists for
that type, in effect reserving it. Hence, over time, pages of the different
types can be clustered together.
When the preferred freelists are expired, the largest possible block is taken
from an alternative list. Buddies that are split from that large block are
placed on the preferred allocation-type freelists to mitigate fragmentation.
This implementation gives best-effort for low fragmentation in all zones.
Ideally, min_free_kbytes needs to be set to a value equal to 4 * (1 <<
(MAX_ORDER-1)) pages in most cases. This would be 16384 on x86 and x86_64 for
example.
Our tests show that about 60-70% of physical memory can be allocated on a
desktop after a few days uptime. In benchmarks and stress tests, we are
finding that 80% of memory is available as contiguous blocks at the end of the
test. To compare, a standard kernel was getting < 1% of memory as large pages
on a desktop and about 8-12% of memory as large pages at the end of stress
tests.
Following this email are 12 patches that implement thie page grouping feature.
The first patch introduces a mechanism for storing flags related to a whole
block of pages. Then allocations are split between movable and all other
allocations. Following that are patches to deal with per-cpu pages and make
the mechanism configurable. The next patch moves free pages between lists
when partially allocated blocks are used for pages of another migrate type.
The second last patch groups reclaimable kernel allocations such as inode
caches together. The final patch related to groupings keeps high-order atomic
allocations.
The last two patches are more concerned with control of fragmentation. The
second last patch biases placement of non-movable allocations towards the
start of memory. This is with a view of supporting memory hot-remove of DIMMs
with higher PFNs in the future. The biasing could be enforced a lot heavier
but it would cost. The last patch agressively clusters reclaimable pages like
inode caches together.
The fragmentation reduction strategy needs to track if pages within a block
can be moved or reclaimed so that pages are freed to the appropriate list.
This patch adds a bitmap for flags affecting a whole a MAX_ORDER block of
pages.
In non-SPARSEMEM configurations, the bitmap is stored in the struct zone and
allocated during initialisation. SPARSEMEM statically allocates the bitmap in
a struct mem_section so that bitmaps do not have to be resized during memory
hotadd. This wastes a small amount of memory per unused section (usually
sizeof(unsigned long)) but the complexity of dynamically allocating the memory
is quite high.
Additional credit to Andy Whitcroft who reviewed up an earlier implementation
of the mechanism an suggested how to make it a *lot* cleaner.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Cc: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:25:47 +04:00
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#include <linux/pageblock-flags.h>
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2013-02-23 04:34:30 +04:00
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#include <linux/page-flags-layout.h>
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2011-07-27 03:09:06 +04:00
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#include <linux/atomic.h>
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[PATCH] Sparsemem build fix
From: Ralf Baechle <ralf@linux-mips.org>
<linux/mmzone.h> uses PAGE_SIZE, PAGE_SHIFT from <asm/page.h> without
including that header itself. For some sparsemem configurations this may
result in build errors like:
CC init/initramfs.o
In file included from include/linux/gfp.h:4,
from include/linux/slab.h:15,
from include/linux/percpu.h:4,
from include/linux/rcupdate.h:41,
from include/linux/dcache.h:10,
from include/linux/fs.h:226,
from init/initramfs.c:2:
include/linux/mmzone.h:498:22: warning: "PAGE_SHIFT" is not defined
In file included from include/linux/gfp.h:4,
from include/linux/slab.h:15,
from include/linux/percpu.h:4,
from include/linux/rcupdate.h:41,
from include/linux/dcache.h:10,
from include/linux/fs.h:226,
from init/initramfs.c:2:
include/linux/mmzone.h:526: error: `PAGE_SIZE' undeclared here (not in a function)
include/linux/mmzone.h: In function `__pfn_to_section':
include/linux/mmzone.h:573: error: `PAGE_SHIFT' undeclared (first use in this function)
include/linux/mmzone.h:573: error: (Each undeclared identifier is reported only once
include/linux/mmzone.h:573: error: for each function it appears in.)
include/linux/mmzone.h: In function `pfn_valid':
include/linux/mmzone.h:578: error: `PAGE_SHIFT' undeclared (first use in this function)
make[1]: *** [init/initramfs.o] Error 1
make: *** [init] Error 2
Signed-off-by: Ralf Baechle <ralf@linux-mips.org>
Seems-reasonable-to: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-06-04 13:51:29 +04:00
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#include <asm/page.h>
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2005-04-17 02:20:36 +04:00
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/* Free memory management - zoned buddy allocator. */
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#ifndef CONFIG_FORCE_MAX_ZONEORDER
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#define MAX_ORDER 11
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#else
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#define MAX_ORDER CONFIG_FORCE_MAX_ZONEORDER
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#endif
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2006-05-21 02:00:31 +04:00
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#define MAX_ORDER_NR_PAGES (1 << (MAX_ORDER - 1))
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2005-04-17 02:20:36 +04:00
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2007-07-17 15:03:16 +04:00
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/*
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* PAGE_ALLOC_COSTLY_ORDER is the order at which allocations are deemed
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* costly to service. That is between allocation orders which should
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2012-04-15 16:42:28 +04:00
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* coalesce naturally under reasonable reclaim pressure and those which
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2007-07-17 15:03:16 +04:00
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* will not.
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*/
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#define PAGE_ALLOC_COSTLY_ORDER 3
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2011-12-29 16:09:50 +04:00
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enum {
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MIGRATE_UNMOVABLE,
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MIGRATE_RECLAIMABLE,
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MIGRATE_MOVABLE,
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MIGRATE_PCPTYPES, /* the number of types on the pcp lists */
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MIGRATE_RESERVE = MIGRATE_PCPTYPES,
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#ifdef CONFIG_CMA
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/*
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* MIGRATE_CMA migration type is designed to mimic the way
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* ZONE_MOVABLE works. Only movable pages can be allocated
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* from MIGRATE_CMA pageblocks and page allocator never
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* implicitly change migration type of MIGRATE_CMA pageblock.
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*
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* The way to use it is to change migratetype of a range of
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* pageblocks to MIGRATE_CMA which can be done by
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* __free_pageblock_cma() function. What is important though
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* is that a range of pageblocks must be aligned to
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* MAX_ORDER_NR_PAGES should biggest page be bigger then
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* a single pageblock.
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*/
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MIGRATE_CMA,
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#endif
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2013-02-23 04:33:58 +04:00
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#ifdef CONFIG_MEMORY_ISOLATION
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2011-12-29 16:09:50 +04:00
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MIGRATE_ISOLATE, /* can't allocate from here */
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2013-02-23 04:33:58 +04:00
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#endif
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2011-12-29 16:09:50 +04:00
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MIGRATE_TYPES
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};
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#ifdef CONFIG_CMA
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# define is_migrate_cma(migratetype) unlikely((migratetype) == MIGRATE_CMA)
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#else
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# define is_migrate_cma(migratetype) false
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#endif
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2007-10-16 12:25:48 +04:00
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#define for_each_migratetype_order(order, type) \
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for (order = 0; order < MAX_ORDER; order++) \
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for (type = 0; type < MIGRATE_TYPES; type++)
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Print out statistics in relation to fragmentation avoidance to /proc/pagetypeinfo
This patch provides fragmentation avoidance statistics via /proc/pagetypeinfo.
The information is collected only on request so there is no runtime overhead.
The statistics are in three parts:
The first part prints information on the size of blocks that pages are
being grouped on and looks like
Page block order: 10
Pages per block: 1024
The second part is a more detailed version of /proc/buddyinfo and looks like
Free pages count per migrate type at order 0 1 2 3 4 5 6 7 8 9 10
Node 0, zone DMA, type Unmovable 0 0 0 0 0 0 0 0 0 0 0
Node 0, zone DMA, type Reclaimable 1 0 0 0 0 0 0 0 0 0 0
Node 0, zone DMA, type Movable 0 0 0 0 0 0 0 0 0 0 0
Node 0, zone DMA, type Reserve 0 4 4 0 0 0 0 1 0 1 0
Node 0, zone Normal, type Unmovable 111 8 4 4 2 3 1 0 0 0 0
Node 0, zone Normal, type Reclaimable 293 89 8 0 0 0 0 0 0 0 0
Node 0, zone Normal, type Movable 1 6 13 9 7 6 3 0 0 0 0
Node 0, zone Normal, type Reserve 0 0 0 0 0 0 0 0 0 0 4
The third part looks like
Number of blocks type Unmovable Reclaimable Movable Reserve
Node 0, zone DMA 0 1 2 1
Node 0, zone Normal 3 17 94 4
To walk the zones within a node with interrupts disabled, walk_zones_in_node()
is introduced and shared between /proc/buddyinfo, /proc/zoneinfo and
/proc/pagetypeinfo to reduce code duplication. It seems specific to what
vmstat.c requires but could be broken out as a general utility function in
mmzone.c if there were other other potential users.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Acked-by: Andy Whitcroft <apw@shadowen.org>
Acked-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:26:02 +04:00
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extern int page_group_by_mobility_disabled;
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static inline int get_pageblock_migratetype(struct page *page)
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{
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return get_pageblock_flags_group(page, PB_migrate, PB_migrate_end);
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}
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2005-04-17 02:20:36 +04:00
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struct free_area {
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2007-10-16 12:25:48 +04:00
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struct list_head free_list[MIGRATE_TYPES];
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2005-04-17 02:20:36 +04:00
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unsigned long nr_free;
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};
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struct pglist_data;
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/*
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* zone->lock and zone->lru_lock are two of the hottest locks in the kernel.
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* So add a wild amount of padding here to ensure that they fall into separate
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* cachelines. There are very few zone structures in the machine, so space
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* consumption is not a concern here.
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*/
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#if defined(CONFIG_SMP)
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struct zone_padding {
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char x[0];
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2006-01-08 12:01:27 +03:00
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} ____cacheline_internodealigned_in_smp;
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2005-04-17 02:20:36 +04:00
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#define ZONE_PADDING(name) struct zone_padding name;
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#else
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#define ZONE_PADDING(name)
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#endif
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2006-06-30 12:55:33 +04:00
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enum zone_stat_item {
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2007-02-10 12:43:02 +03:00
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/* First 128 byte cacheline (assuming 64 bit words) */
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2007-02-10 12:43:02 +03:00
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NR_FREE_PAGES,
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mm: page_alloc: fair zone allocator policy
Each zone that holds userspace pages of one workload must be aged at a
speed proportional to the zone size. Otherwise, the time an individual
page gets to stay in memory depends on the zone it happened to be
allocated in. Asymmetry in the zone aging creates rather unpredictable
aging behavior and results in the wrong pages being reclaimed, activated
etc.
But exactly this happens right now because of the way the page allocator
and kswapd interact. The page allocator uses per-node lists of all zones
in the system, ordered by preference, when allocating a new page. When
the first iteration does not yield any results, kswapd is woken up and the
allocator retries. Due to the way kswapd reclaims zones below the high
watermark while a zone can be allocated from when it is above the low
watermark, the allocator may keep kswapd running while kswapd reclaim
ensures that the page allocator can keep allocating from the first zone in
the zonelist for extended periods of time. Meanwhile the other zones
rarely see new allocations and thus get aged much slower in comparison.
The result is that the occasional page placed in lower zones gets
relatively more time in memory, even gets promoted to the active list
after its peers have long been evicted. Meanwhile, the bulk of the
working set may be thrashing on the preferred zone even though there may
be significant amounts of memory available in the lower zones.
Even the most basic test -- repeatedly reading a file slightly bigger than
memory -- shows how broken the zone aging is. In this scenario, no single
page should be able stay in memory long enough to get referenced twice and
activated, but activation happens in spades:
$ grep active_file /proc/zoneinfo
nr_inactive_file 0
nr_active_file 0
nr_inactive_file 0
nr_active_file 8
nr_inactive_file 1582
nr_active_file 11994
$ cat data data data data >/dev/null
$ grep active_file /proc/zoneinfo
nr_inactive_file 0
nr_active_file 70
nr_inactive_file 258753
nr_active_file 443214
nr_inactive_file 149793
nr_active_file 12021
Fix this with a very simple round robin allocator. Each zone is allowed a
batch of allocations that is proportional to the zone's size, after which
it is treated as full. The batch counters are reset when all zones have
been tried and the allocator enters the slowpath and kicks off kswapd
reclaim. Allocation and reclaim is now fairly spread out to all
available/allowable zones:
$ grep active_file /proc/zoneinfo
nr_inactive_file 0
nr_active_file 0
nr_inactive_file 174
nr_active_file 4865
nr_inactive_file 53
nr_active_file 860
$ cat data data data data >/dev/null
$ grep active_file /proc/zoneinfo
nr_inactive_file 0
nr_active_file 0
nr_inactive_file 666622
nr_active_file 4988
nr_inactive_file 190969
nr_active_file 937
When zone_reclaim_mode is enabled, allocations will now spread out to all
zones on the local node, not just the first preferred zone (which on a 4G
node might be a tiny Normal zone).
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: Mel Gorman <mgorman@suse.de>
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Paul Bolle <paul.bollee@gmail.com>
Cc: Zlatko Calusic <zcalusic@bitsync.net>
Tested-by: Kevin Hilman <khilman@linaro.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-09-12 01:20:47 +04:00
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NR_ALLOC_BATCH,
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2008-10-19 07:26:14 +04:00
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NR_LRU_BASE,
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2008-10-19 07:26:32 +04:00
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NR_INACTIVE_ANON = NR_LRU_BASE, /* must match order of LRU_[IN]ACTIVE */
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NR_ACTIVE_ANON, /* " " " " " */
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NR_INACTIVE_FILE, /* " " " " " */
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NR_ACTIVE_FILE, /* " " " " " */
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Unevictable LRU Infrastructure
When the system contains lots of mlocked or otherwise unevictable pages,
the pageout code (kswapd) can spend lots of time scanning over these
pages. Worse still, the presence of lots of unevictable pages can confuse
kswapd into thinking that more aggressive pageout modes are required,
resulting in all kinds of bad behaviour.
Infrastructure to manage pages excluded from reclaim--i.e., hidden from
vmscan. Based on a patch by Larry Woodman of Red Hat. Reworked to
maintain "unevictable" pages on a separate per-zone LRU list, to "hide"
them from vmscan.
Kosaki Motohiro added the support for the memory controller unevictable
lru list.
Pages on the unevictable list have both PG_unevictable and PG_lru set.
Thus, PG_unevictable is analogous to and mutually exclusive with
PG_active--it specifies which LRU list the page is on.
The unevictable infrastructure is enabled by a new mm Kconfig option
[CONFIG_]UNEVICTABLE_LRU.
A new function 'page_evictable(page, vma)' in vmscan.c tests whether or
not a page may be evictable. Subsequent patches will add the various
!evictable tests. We'll want to keep these tests light-weight for use in
shrink_active_list() and, possibly, the fault path.
To avoid races between tasks putting pages [back] onto an LRU list and
tasks that might be moving the page from non-evictable to evictable state,
the new function 'putback_lru_page()' -- inverse to 'isolate_lru_page()'
-- tests the "evictability" of a page after placing it on the LRU, before
dropping the reference. If the page has become unevictable,
putback_lru_page() will redo the 'putback', thus moving the page to the
unevictable list. This way, we avoid "stranding" evictable pages on the
unevictable list.
[akpm@linux-foundation.org: fix fallout from out-of-order merge]
[riel@redhat.com: fix UNEVICTABLE_LRU and !PROC_PAGE_MONITOR build]
[nishimura@mxp.nes.nec.co.jp: remove redundant mapping check]
[kosaki.motohiro@jp.fujitsu.com: unevictable-lru-infrastructure: putback_lru_page()/unevictable page handling rework]
[kosaki.motohiro@jp.fujitsu.com: kill unnecessary lock_page() in vmscan.c]
[kosaki.motohiro@jp.fujitsu.com: revert migration change of unevictable lru infrastructure]
[kosaki.motohiro@jp.fujitsu.com: revert to unevictable-lru-infrastructure-kconfig-fix.patch]
[kosaki.motohiro@jp.fujitsu.com: restore patch failure of vmstat-unevictable-and-mlocked-pages-vm-events.patch]
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Rik van Riel <riel@redhat.com>
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Debugged-by: Benjamin Kidwell <benjkidwell@yahoo.com>
Signed-off-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 07:26:39 +04:00
|
|
|
NR_UNEVICTABLE, /* " " " " " */
|
2008-10-19 07:26:51 +04:00
|
|
|
NR_MLOCK, /* mlock()ed pages found and moved off LRU */
|
2006-06-30 12:55:36 +04:00
|
|
|
NR_ANON_PAGES, /* Mapped anonymous pages */
|
|
|
|
NR_FILE_MAPPED, /* pagecache pages mapped into pagetables.
|
2006-06-30 12:55:34 +04:00
|
|
|
only modified from process context */
|
2006-06-30 12:55:35 +04:00
|
|
|
NR_FILE_PAGES,
|
2006-06-30 12:55:39 +04:00
|
|
|
NR_FILE_DIRTY,
|
2006-06-30 12:55:40 +04:00
|
|
|
NR_WRITEBACK,
|
2007-02-10 12:43:02 +03:00
|
|
|
NR_SLAB_RECLAIMABLE,
|
|
|
|
NR_SLAB_UNRECLAIMABLE,
|
|
|
|
NR_PAGETABLE, /* used for pagetables */
|
2009-09-22 04:01:32 +04:00
|
|
|
NR_KERNEL_STACK,
|
|
|
|
/* Second 128 byte cacheline */
|
2006-06-30 12:55:40 +04:00
|
|
|
NR_UNSTABLE_NFS, /* NFS unstable pages */
|
2006-06-30 12:55:41 +04:00
|
|
|
NR_BOUNCE,
|
2006-09-27 12:50:00 +04:00
|
|
|
NR_VMSCAN_WRITE,
|
2011-11-01 04:07:59 +04:00
|
|
|
NR_VMSCAN_IMMEDIATE, /* Prioritise for reclaim when writeback ends */
|
2008-04-30 11:54:38 +04:00
|
|
|
NR_WRITEBACK_TEMP, /* Writeback using temporary buffers */
|
2009-09-22 04:01:37 +04:00
|
|
|
NR_ISOLATED_ANON, /* Temporary isolated pages from anon lru */
|
|
|
|
NR_ISOLATED_FILE, /* Temporary isolated pages from file lru */
|
2009-09-22 04:01:33 +04:00
|
|
|
NR_SHMEM, /* shmem pages (included tmpfs/GEM pages) */
|
2010-10-27 01:21:35 +04:00
|
|
|
NR_DIRTIED, /* page dirtyings since bootup */
|
|
|
|
NR_WRITTEN, /* page writings since bootup */
|
2006-06-30 12:55:44 +04:00
|
|
|
#ifdef CONFIG_NUMA
|
|
|
|
NUMA_HIT, /* allocated in intended node */
|
|
|
|
NUMA_MISS, /* allocated in non intended node */
|
|
|
|
NUMA_FOREIGN, /* was intended here, hit elsewhere */
|
|
|
|
NUMA_INTERLEAVE_HIT, /* interleaver preferred this zone */
|
|
|
|
NUMA_LOCAL, /* allocation from local node */
|
|
|
|
NUMA_OTHER, /* allocation from other node */
|
|
|
|
#endif
|
2011-01-14 02:46:58 +03:00
|
|
|
NR_ANON_TRANSPARENT_HUGEPAGES,
|
2012-10-09 03:32:02 +04:00
|
|
|
NR_FREE_CMA_PAGES,
|
2006-06-30 12:55:33 +04:00
|
|
|
NR_VM_ZONE_STAT_ITEMS };
|
|
|
|
|
2008-10-19 07:26:32 +04:00
|
|
|
/*
|
|
|
|
* We do arithmetic on the LRU lists in various places in the code,
|
|
|
|
* so it is important to keep the active lists LRU_ACTIVE higher in
|
|
|
|
* the array than the corresponding inactive lists, and to keep
|
|
|
|
* the *_FILE lists LRU_FILE higher than the corresponding _ANON lists.
|
|
|
|
*
|
|
|
|
* This has to be kept in sync with the statistics in zone_stat_item
|
|
|
|
* above and the descriptions in vmstat_text in mm/vmstat.c
|
|
|
|
*/
|
|
|
|
#define LRU_BASE 0
|
|
|
|
#define LRU_ACTIVE 1
|
|
|
|
#define LRU_FILE 2
|
|
|
|
|
2008-10-19 07:26:14 +04:00
|
|
|
enum lru_list {
|
2008-10-19 07:26:32 +04:00
|
|
|
LRU_INACTIVE_ANON = LRU_BASE,
|
|
|
|
LRU_ACTIVE_ANON = LRU_BASE + LRU_ACTIVE,
|
|
|
|
LRU_INACTIVE_FILE = LRU_BASE + LRU_FILE,
|
|
|
|
LRU_ACTIVE_FILE = LRU_BASE + LRU_FILE + LRU_ACTIVE,
|
Unevictable LRU Infrastructure
When the system contains lots of mlocked or otherwise unevictable pages,
the pageout code (kswapd) can spend lots of time scanning over these
pages. Worse still, the presence of lots of unevictable pages can confuse
kswapd into thinking that more aggressive pageout modes are required,
resulting in all kinds of bad behaviour.
Infrastructure to manage pages excluded from reclaim--i.e., hidden from
vmscan. Based on a patch by Larry Woodman of Red Hat. Reworked to
maintain "unevictable" pages on a separate per-zone LRU list, to "hide"
them from vmscan.
Kosaki Motohiro added the support for the memory controller unevictable
lru list.
Pages on the unevictable list have both PG_unevictable and PG_lru set.
Thus, PG_unevictable is analogous to and mutually exclusive with
PG_active--it specifies which LRU list the page is on.
The unevictable infrastructure is enabled by a new mm Kconfig option
[CONFIG_]UNEVICTABLE_LRU.
A new function 'page_evictable(page, vma)' in vmscan.c tests whether or
not a page may be evictable. Subsequent patches will add the various
!evictable tests. We'll want to keep these tests light-weight for use in
shrink_active_list() and, possibly, the fault path.
To avoid races between tasks putting pages [back] onto an LRU list and
tasks that might be moving the page from non-evictable to evictable state,
the new function 'putback_lru_page()' -- inverse to 'isolate_lru_page()'
-- tests the "evictability" of a page after placing it on the LRU, before
dropping the reference. If the page has become unevictable,
putback_lru_page() will redo the 'putback', thus moving the page to the
unevictable list. This way, we avoid "stranding" evictable pages on the
unevictable list.
[akpm@linux-foundation.org: fix fallout from out-of-order merge]
[riel@redhat.com: fix UNEVICTABLE_LRU and !PROC_PAGE_MONITOR build]
[nishimura@mxp.nes.nec.co.jp: remove redundant mapping check]
[kosaki.motohiro@jp.fujitsu.com: unevictable-lru-infrastructure: putback_lru_page()/unevictable page handling rework]
[kosaki.motohiro@jp.fujitsu.com: kill unnecessary lock_page() in vmscan.c]
[kosaki.motohiro@jp.fujitsu.com: revert migration change of unevictable lru infrastructure]
[kosaki.motohiro@jp.fujitsu.com: revert to unevictable-lru-infrastructure-kconfig-fix.patch]
[kosaki.motohiro@jp.fujitsu.com: restore patch failure of vmstat-unevictable-and-mlocked-pages-vm-events.patch]
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Rik van Riel <riel@redhat.com>
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Debugged-by: Benjamin Kidwell <benjkidwell@yahoo.com>
Signed-off-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 07:26:39 +04:00
|
|
|
LRU_UNEVICTABLE,
|
|
|
|
NR_LRU_LISTS
|
|
|
|
};
|
2008-10-19 07:26:14 +04:00
|
|
|
|
2012-01-13 05:20:01 +04:00
|
|
|
#define for_each_lru(lru) for (lru = 0; lru < NR_LRU_LISTS; lru++)
|
2008-10-19 07:26:14 +04:00
|
|
|
|
2012-01-13 05:20:01 +04:00
|
|
|
#define for_each_evictable_lru(lru) for (lru = 0; lru <= LRU_ACTIVE_FILE; lru++)
|
Unevictable LRU Infrastructure
When the system contains lots of mlocked or otherwise unevictable pages,
the pageout code (kswapd) can spend lots of time scanning over these
pages. Worse still, the presence of lots of unevictable pages can confuse
kswapd into thinking that more aggressive pageout modes are required,
resulting in all kinds of bad behaviour.
Infrastructure to manage pages excluded from reclaim--i.e., hidden from
vmscan. Based on a patch by Larry Woodman of Red Hat. Reworked to
maintain "unevictable" pages on a separate per-zone LRU list, to "hide"
them from vmscan.
Kosaki Motohiro added the support for the memory controller unevictable
lru list.
Pages on the unevictable list have both PG_unevictable and PG_lru set.
Thus, PG_unevictable is analogous to and mutually exclusive with
PG_active--it specifies which LRU list the page is on.
The unevictable infrastructure is enabled by a new mm Kconfig option
[CONFIG_]UNEVICTABLE_LRU.
A new function 'page_evictable(page, vma)' in vmscan.c tests whether or
not a page may be evictable. Subsequent patches will add the various
!evictable tests. We'll want to keep these tests light-weight for use in
shrink_active_list() and, possibly, the fault path.
To avoid races between tasks putting pages [back] onto an LRU list and
tasks that might be moving the page from non-evictable to evictable state,
the new function 'putback_lru_page()' -- inverse to 'isolate_lru_page()'
-- tests the "evictability" of a page after placing it on the LRU, before
dropping the reference. If the page has become unevictable,
putback_lru_page() will redo the 'putback', thus moving the page to the
unevictable list. This way, we avoid "stranding" evictable pages on the
unevictable list.
[akpm@linux-foundation.org: fix fallout from out-of-order merge]
[riel@redhat.com: fix UNEVICTABLE_LRU and !PROC_PAGE_MONITOR build]
[nishimura@mxp.nes.nec.co.jp: remove redundant mapping check]
[kosaki.motohiro@jp.fujitsu.com: unevictable-lru-infrastructure: putback_lru_page()/unevictable page handling rework]
[kosaki.motohiro@jp.fujitsu.com: kill unnecessary lock_page() in vmscan.c]
[kosaki.motohiro@jp.fujitsu.com: revert migration change of unevictable lru infrastructure]
[kosaki.motohiro@jp.fujitsu.com: revert to unevictable-lru-infrastructure-kconfig-fix.patch]
[kosaki.motohiro@jp.fujitsu.com: restore patch failure of vmstat-unevictable-and-mlocked-pages-vm-events.patch]
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Rik van Riel <riel@redhat.com>
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Debugged-by: Benjamin Kidwell <benjkidwell@yahoo.com>
Signed-off-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 07:26:39 +04:00
|
|
|
|
2012-01-13 05:20:01 +04:00
|
|
|
static inline int is_file_lru(enum lru_list lru)
|
2008-10-19 07:26:32 +04:00
|
|
|
{
|
2012-01-13 05:20:01 +04:00
|
|
|
return (lru == LRU_INACTIVE_FILE || lru == LRU_ACTIVE_FILE);
|
2008-10-19 07:26:32 +04:00
|
|
|
}
|
|
|
|
|
2012-01-13 05:20:01 +04:00
|
|
|
static inline int is_active_lru(enum lru_list lru)
|
2008-10-19 07:26:14 +04:00
|
|
|
{
|
2012-01-13 05:20:01 +04:00
|
|
|
return (lru == LRU_ACTIVE_ANON || lru == LRU_ACTIVE_FILE);
|
2008-10-19 07:26:14 +04:00
|
|
|
}
|
|
|
|
|
2012-01-13 05:20:01 +04:00
|
|
|
static inline int is_unevictable_lru(enum lru_list lru)
|
Unevictable LRU Infrastructure
When the system contains lots of mlocked or otherwise unevictable pages,
the pageout code (kswapd) can spend lots of time scanning over these
pages. Worse still, the presence of lots of unevictable pages can confuse
kswapd into thinking that more aggressive pageout modes are required,
resulting in all kinds of bad behaviour.
Infrastructure to manage pages excluded from reclaim--i.e., hidden from
vmscan. Based on a patch by Larry Woodman of Red Hat. Reworked to
maintain "unevictable" pages on a separate per-zone LRU list, to "hide"
them from vmscan.
Kosaki Motohiro added the support for the memory controller unevictable
lru list.
Pages on the unevictable list have both PG_unevictable and PG_lru set.
Thus, PG_unevictable is analogous to and mutually exclusive with
PG_active--it specifies which LRU list the page is on.
The unevictable infrastructure is enabled by a new mm Kconfig option
[CONFIG_]UNEVICTABLE_LRU.
A new function 'page_evictable(page, vma)' in vmscan.c tests whether or
not a page may be evictable. Subsequent patches will add the various
!evictable tests. We'll want to keep these tests light-weight for use in
shrink_active_list() and, possibly, the fault path.
To avoid races between tasks putting pages [back] onto an LRU list and
tasks that might be moving the page from non-evictable to evictable state,
the new function 'putback_lru_page()' -- inverse to 'isolate_lru_page()'
-- tests the "evictability" of a page after placing it on the LRU, before
dropping the reference. If the page has become unevictable,
putback_lru_page() will redo the 'putback', thus moving the page to the
unevictable list. This way, we avoid "stranding" evictable pages on the
unevictable list.
[akpm@linux-foundation.org: fix fallout from out-of-order merge]
[riel@redhat.com: fix UNEVICTABLE_LRU and !PROC_PAGE_MONITOR build]
[nishimura@mxp.nes.nec.co.jp: remove redundant mapping check]
[kosaki.motohiro@jp.fujitsu.com: unevictable-lru-infrastructure: putback_lru_page()/unevictable page handling rework]
[kosaki.motohiro@jp.fujitsu.com: kill unnecessary lock_page() in vmscan.c]
[kosaki.motohiro@jp.fujitsu.com: revert migration change of unevictable lru infrastructure]
[kosaki.motohiro@jp.fujitsu.com: revert to unevictable-lru-infrastructure-kconfig-fix.patch]
[kosaki.motohiro@jp.fujitsu.com: restore patch failure of vmstat-unevictable-and-mlocked-pages-vm-events.patch]
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Rik van Riel <riel@redhat.com>
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Debugged-by: Benjamin Kidwell <benjkidwell@yahoo.com>
Signed-off-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 07:26:39 +04:00
|
|
|
{
|
2012-01-13 05:20:01 +04:00
|
|
|
return (lru == LRU_UNEVICTABLE);
|
Unevictable LRU Infrastructure
When the system contains lots of mlocked or otherwise unevictable pages,
the pageout code (kswapd) can spend lots of time scanning over these
pages. Worse still, the presence of lots of unevictable pages can confuse
kswapd into thinking that more aggressive pageout modes are required,
resulting in all kinds of bad behaviour.
Infrastructure to manage pages excluded from reclaim--i.e., hidden from
vmscan. Based on a patch by Larry Woodman of Red Hat. Reworked to
maintain "unevictable" pages on a separate per-zone LRU list, to "hide"
them from vmscan.
Kosaki Motohiro added the support for the memory controller unevictable
lru list.
Pages on the unevictable list have both PG_unevictable and PG_lru set.
Thus, PG_unevictable is analogous to and mutually exclusive with
PG_active--it specifies which LRU list the page is on.
The unevictable infrastructure is enabled by a new mm Kconfig option
[CONFIG_]UNEVICTABLE_LRU.
A new function 'page_evictable(page, vma)' in vmscan.c tests whether or
not a page may be evictable. Subsequent patches will add the various
!evictable tests. We'll want to keep these tests light-weight for use in
shrink_active_list() and, possibly, the fault path.
To avoid races between tasks putting pages [back] onto an LRU list and
tasks that might be moving the page from non-evictable to evictable state,
the new function 'putback_lru_page()' -- inverse to 'isolate_lru_page()'
-- tests the "evictability" of a page after placing it on the LRU, before
dropping the reference. If the page has become unevictable,
putback_lru_page() will redo the 'putback', thus moving the page to the
unevictable list. This way, we avoid "stranding" evictable pages on the
unevictable list.
[akpm@linux-foundation.org: fix fallout from out-of-order merge]
[riel@redhat.com: fix UNEVICTABLE_LRU and !PROC_PAGE_MONITOR build]
[nishimura@mxp.nes.nec.co.jp: remove redundant mapping check]
[kosaki.motohiro@jp.fujitsu.com: unevictable-lru-infrastructure: putback_lru_page()/unevictable page handling rework]
[kosaki.motohiro@jp.fujitsu.com: kill unnecessary lock_page() in vmscan.c]
[kosaki.motohiro@jp.fujitsu.com: revert migration change of unevictable lru infrastructure]
[kosaki.motohiro@jp.fujitsu.com: revert to unevictable-lru-infrastructure-kconfig-fix.patch]
[kosaki.motohiro@jp.fujitsu.com: restore patch failure of vmstat-unevictable-and-mlocked-pages-vm-events.patch]
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Rik van Riel <riel@redhat.com>
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Debugged-by: Benjamin Kidwell <benjkidwell@yahoo.com>
Signed-off-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 07:26:39 +04:00
|
|
|
}
|
|
|
|
|
2012-05-30 02:06:53 +04:00
|
|
|
struct zone_reclaim_stat {
|
|
|
|
/*
|
|
|
|
* The pageout code in vmscan.c keeps track of how many of the
|
2012-06-29 16:45:58 +04:00
|
|
|
* mem/swap backed and file backed pages are referenced.
|
2012-05-30 02:06:53 +04:00
|
|
|
* The higher the rotated/scanned ratio, the more valuable
|
|
|
|
* that cache is.
|
|
|
|
*
|
|
|
|
* The anon LRU stats live in [0], file LRU stats in [1]
|
|
|
|
*/
|
|
|
|
unsigned long recent_rotated[2];
|
|
|
|
unsigned long recent_scanned[2];
|
|
|
|
};
|
|
|
|
|
2012-01-13 05:18:10 +04:00
|
|
|
struct lruvec {
|
|
|
|
struct list_head lists[NR_LRU_LISTS];
|
2012-05-30 02:06:53 +04:00
|
|
|
struct zone_reclaim_stat reclaim_stat;
|
2012-08-01 03:43:02 +04:00
|
|
|
#ifdef CONFIG_MEMCG
|
2012-05-30 02:06:58 +04:00
|
|
|
struct zone *zone;
|
|
|
|
#endif
|
2012-01-13 05:18:10 +04:00
|
|
|
};
|
|
|
|
|
memcg: consolidate memory cgroup lru stat functions
In mm/memcontrol.c, there are many lru stat functions as..
mem_cgroup_zone_nr_lru_pages
mem_cgroup_node_nr_file_lru_pages
mem_cgroup_nr_file_lru_pages
mem_cgroup_node_nr_anon_lru_pages
mem_cgroup_nr_anon_lru_pages
mem_cgroup_node_nr_unevictable_lru_pages
mem_cgroup_nr_unevictable_lru_pages
mem_cgroup_node_nr_lru_pages
mem_cgroup_nr_lru_pages
mem_cgroup_get_local_zonestat
Some of them are under #ifdef MAX_NUMNODES >1 and others are not.
This seems bad. This patch consolidates all functions into
mem_cgroup_zone_nr_lru_pages()
mem_cgroup_node_nr_lru_pages()
mem_cgroup_nr_lru_pages()
For these functions, "which LRU?" information is passed by a mask.
example:
mem_cgroup_nr_lru_pages(mem, BIT(LRU_ACTIVE_ANON))
And I added some macro as ALL_LRU, ALL_LRU_FILE, ALL_LRU_ANON.
example:
mem_cgroup_nr_lru_pages(mem, ALL_LRU)
BTW, considering layout of NUMA memory placement of counters, this patch seems
to be better.
Now, when we gather all LRU information, we scan in following orer
for_each_lru -> for_each_node -> for_each_zone.
This means we'll touch cache lines in different node in turn.
After patch, we'll scan
for_each_node -> for_each_zone -> for_each_lru(mask)
Then, we'll gather information in the same cacheline at once.
[akpm@linux-foundation.org: fix warnigns, build error]
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Cc: Balbir Singh <bsingharora@gmail.com>
Cc: Michal Hocko <mhocko@suse.cz>
Cc: Ying Han <yinghan@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-07-27 03:08:22 +04:00
|
|
|
/* Mask used at gathering information at once (see memcontrol.c) */
|
|
|
|
#define LRU_ALL_FILE (BIT(LRU_INACTIVE_FILE) | BIT(LRU_ACTIVE_FILE))
|
|
|
|
#define LRU_ALL_ANON (BIT(LRU_INACTIVE_ANON) | BIT(LRU_ACTIVE_ANON))
|
|
|
|
#define LRU_ALL ((1 << NR_LRU_LISTS) - 1)
|
|
|
|
|
2011-11-01 04:06:51 +04:00
|
|
|
/* Isolate clean file */
|
2012-05-30 02:06:54 +04:00
|
|
|
#define ISOLATE_CLEAN ((__force isolate_mode_t)0x1)
|
2011-11-01 04:06:55 +04:00
|
|
|
/* Isolate unmapped file */
|
2012-05-30 02:06:54 +04:00
|
|
|
#define ISOLATE_UNMAPPED ((__force isolate_mode_t)0x2)
|
2012-01-13 05:19:38 +04:00
|
|
|
/* Isolate for asynchronous migration */
|
2012-05-30 02:06:54 +04:00
|
|
|
#define ISOLATE_ASYNC_MIGRATE ((__force isolate_mode_t)0x4)
|
2012-10-09 03:33:48 +04:00
|
|
|
/* Isolate unevictable pages */
|
|
|
|
#define ISOLATE_UNEVICTABLE ((__force isolate_mode_t)0x8)
|
2011-11-01 04:06:47 +04:00
|
|
|
|
|
|
|
/* LRU Isolation modes. */
|
|
|
|
typedef unsigned __bitwise__ isolate_mode_t;
|
|
|
|
|
2009-06-17 02:32:12 +04:00
|
|
|
enum zone_watermarks {
|
|
|
|
WMARK_MIN,
|
|
|
|
WMARK_LOW,
|
|
|
|
WMARK_HIGH,
|
|
|
|
NR_WMARK
|
|
|
|
};
|
|
|
|
|
|
|
|
#define min_wmark_pages(z) (z->watermark[WMARK_MIN])
|
|
|
|
#define low_wmark_pages(z) (z->watermark[WMARK_LOW])
|
|
|
|
#define high_wmark_pages(z) (z->watermark[WMARK_HIGH])
|
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
struct per_cpu_pages {
|
|
|
|
int count; /* number of pages in the list */
|
|
|
|
int high; /* high watermark, emptying needed */
|
|
|
|
int batch; /* chunk size for buddy add/remove */
|
page-allocator: split per-cpu list into one-list-per-migrate-type
The following two patches remove searching in the page allocator fast-path
by maintaining multiple free-lists in the per-cpu structure. At the time
the search was introduced, increasing the per-cpu structures would waste a
lot of memory as per-cpu structures were statically allocated at
compile-time. This is no longer the case.
The patches are as follows. They are based on mmotm-2009-08-27.
Patch 1 adds multiple lists to struct per_cpu_pages, one per
migratetype that can be stored on the PCP lists.
Patch 2 notes that the pcpu drain path check empty lists multiple times. The
patch reduces the number of checks by maintaining a count of free
lists encountered. Lists containing pages will then free multiple
pages in batch
The patches were tested with kernbench, netperf udp/tcp, hackbench and
sysbench. The netperf tests were not bound to any CPU in particular and
were run such that the results should be 99% confidence that the reported
results are within 1% of the estimated mean. sysbench was run with a
postgres background and read-only tests. Similar to netperf, it was run
multiple times so that it's 99% confidence results are within 1%. The
patches were tested on x86, x86-64 and ppc64 as
x86: Intel Pentium D 3GHz with 8G RAM (no-brand machine)
kernbench - No significant difference, variance well within noise
netperf-udp - 1.34% to 2.28% gain
netperf-tcp - 0.45% to 1.22% gain
hackbench - Small variances, very close to noise
sysbench - Very small gains
x86-64: AMD Phenom 9950 1.3GHz with 8G RAM (no-brand machine)
kernbench - No significant difference, variance well within noise
netperf-udp - 1.83% to 10.42% gains
netperf-tcp - No conclusive until buffer >= PAGE_SIZE
4096 +15.83%
8192 + 0.34% (not significant)
16384 + 1%
hackbench - Small gains, very close to noise
sysbench - 0.79% to 1.6% gain
ppc64: PPC970MP 2.5GHz with 10GB RAM (it's a terrasoft powerstation)
kernbench - No significant difference, variance well within noise
netperf-udp - 2-3% gain for almost all buffer sizes tested
netperf-tcp - losses on small buffers, gains on larger buffers
possibly indicates some bad caching effect.
hackbench - No significant difference
sysbench - 2-4% gain
This patch:
Currently the per-cpu page allocator searches the PCP list for pages of
the correct migrate-type to reduce the possibility of pages being
inappropriate placed from a fragmentation perspective. This search is
potentially expensive in a fast-path and undesirable. Splitting the
per-cpu list into multiple lists increases the size of a per-cpu structure
and this was potentially a major problem at the time the search was
introduced. These problem has been mitigated as now only the necessary
number of structures is allocated for the running system.
This patch replaces a list search in the per-cpu allocator with one list
per migrate type. The potential snag with this approach is when bulk
freeing pages. We round-robin free pages based on migrate type which has
little bearing on the cache hotness of the page and potentially checks
empty lists repeatedly in the event the majority of PCP pages are of one
type.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Acked-by: Nick Piggin <npiggin@suse.de>
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Minchan Kim <minchan.kim@gmail.com>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-22 04:03:19 +04:00
|
|
|
|
|
|
|
/* Lists of pages, one per migrate type stored on the pcp-lists */
|
|
|
|
struct list_head lists[MIGRATE_PCPTYPES];
|
2005-04-17 02:20:36 +04:00
|
|
|
};
|
|
|
|
|
|
|
|
struct per_cpu_pageset {
|
2008-02-05 09:29:19 +03:00
|
|
|
struct per_cpu_pages pcp;
|
2007-05-09 13:35:14 +04:00
|
|
|
#ifdef CONFIG_NUMA
|
|
|
|
s8 expire;
|
|
|
|
#endif
|
2006-06-30 12:55:33 +04:00
|
|
|
#ifdef CONFIG_SMP
|
2006-09-01 08:27:35 +04:00
|
|
|
s8 stat_threshold;
|
2006-06-30 12:55:33 +04:00
|
|
|
s8 vm_stat_diff[NR_VM_ZONE_STAT_ITEMS];
|
|
|
|
#endif
|
2010-01-05 09:34:51 +03:00
|
|
|
};
|
2005-06-22 04:14:47 +04:00
|
|
|
|
2008-04-28 13:12:54 +04:00
|
|
|
#endif /* !__GENERATING_BOUNDS.H */
|
|
|
|
|
2006-09-26 10:31:13 +04:00
|
|
|
enum zone_type {
|
2007-02-10 12:43:10 +03:00
|
|
|
#ifdef CONFIG_ZONE_DMA
|
2006-09-26 10:31:13 +04:00
|
|
|
/*
|
|
|
|
* ZONE_DMA is used when there are devices that are not able
|
|
|
|
* to do DMA to all of addressable memory (ZONE_NORMAL). Then we
|
|
|
|
* carve out the portion of memory that is needed for these devices.
|
|
|
|
* The range is arch specific.
|
|
|
|
*
|
|
|
|
* Some examples
|
|
|
|
*
|
|
|
|
* Architecture Limit
|
|
|
|
* ---------------------------
|
|
|
|
* parisc, ia64, sparc <4G
|
|
|
|
* s390 <2G
|
|
|
|
* arm Various
|
|
|
|
* alpha Unlimited or 0-16MB.
|
|
|
|
*
|
|
|
|
* i386, x86_64 and multiple other arches
|
|
|
|
* <16M.
|
|
|
|
*/
|
|
|
|
ZONE_DMA,
|
2007-02-10 12:43:10 +03:00
|
|
|
#endif
|
2006-09-26 10:31:13 +04:00
|
|
|
#ifdef CONFIG_ZONE_DMA32
|
2006-09-26 10:31:13 +04:00
|
|
|
/*
|
|
|
|
* x86_64 needs two ZONE_DMAs because it supports devices that are
|
|
|
|
* only able to do DMA to the lower 16M but also 32 bit devices that
|
|
|
|
* can only do DMA areas below 4G.
|
|
|
|
*/
|
|
|
|
ZONE_DMA32,
|
2006-09-26 10:31:13 +04:00
|
|
|
#endif
|
2006-09-26 10:31:13 +04:00
|
|
|
/*
|
|
|
|
* Normal addressable memory is in ZONE_NORMAL. DMA operations can be
|
|
|
|
* performed on pages in ZONE_NORMAL if the DMA devices support
|
|
|
|
* transfers to all addressable memory.
|
|
|
|
*/
|
|
|
|
ZONE_NORMAL,
|
2006-09-26 10:31:14 +04:00
|
|
|
#ifdef CONFIG_HIGHMEM
|
2006-09-26 10:31:13 +04:00
|
|
|
/*
|
|
|
|
* A memory area that is only addressable by the kernel through
|
|
|
|
* mapping portions into its own address space. This is for example
|
|
|
|
* used by i386 to allow the kernel to address the memory beyond
|
|
|
|
* 900MB. The kernel will set up special mappings (page
|
|
|
|
* table entries on i386) for each page that the kernel needs to
|
|
|
|
* access.
|
|
|
|
*/
|
|
|
|
ZONE_HIGHMEM,
|
2006-09-26 10:31:14 +04:00
|
|
|
#endif
|
2007-07-17 15:03:12 +04:00
|
|
|
ZONE_MOVABLE,
|
2008-04-28 13:12:54 +04:00
|
|
|
__MAX_NR_ZONES
|
2006-09-26 10:31:13 +04:00
|
|
|
};
|
2005-04-17 02:20:36 +04:00
|
|
|
|
2008-04-28 13:12:54 +04:00
|
|
|
#ifndef __GENERATING_BOUNDS_H
|
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
struct zone {
|
|
|
|
/* Fields commonly accessed by the page allocator */
|
2009-06-17 02:32:12 +04:00
|
|
|
|
|
|
|
/* zone watermarks, access with *_wmark_pages(zone) macros */
|
|
|
|
unsigned long watermark[NR_WMARK];
|
|
|
|
|
2010-09-10 03:38:17 +04:00
|
|
|
/*
|
|
|
|
* When free pages are below this point, additional steps are taken
|
|
|
|
* when reading the number of free pages to avoid per-cpu counter
|
|
|
|
* drift allowing watermarks to be breached
|
|
|
|
*/
|
|
|
|
unsigned long percpu_drift_mark;
|
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
/*
|
|
|
|
* We don't know if the memory that we're going to allocate will be freeable
|
|
|
|
* or/and it will be released eventually, so to avoid totally wasting several
|
|
|
|
* GB of ram we must reserve some of the lower zone memory (otherwise we risk
|
|
|
|
* to run OOM on the lower zones despite there's tons of freeable ram
|
|
|
|
* on the higher zones). This array is recalculated at runtime if the
|
|
|
|
* sysctl_lowmem_reserve_ratio sysctl changes.
|
|
|
|
*/
|
|
|
|
unsigned long lowmem_reserve[MAX_NR_ZONES];
|
|
|
|
|
2012-01-11 03:07:42 +04:00
|
|
|
/*
|
|
|
|
* This is a per-zone reserve of pages that should not be
|
|
|
|
* considered dirtyable memory.
|
|
|
|
*/
|
|
|
|
unsigned long dirty_balance_reserve;
|
|
|
|
|
2005-06-22 04:14:47 +04:00
|
|
|
#ifdef CONFIG_NUMA
|
2006-09-27 12:50:08 +04:00
|
|
|
int node;
|
2006-07-03 11:24:13 +04:00
|
|
|
/*
|
|
|
|
* zone reclaim becomes active if more unmapped pages exist.
|
|
|
|
*/
|
2006-09-26 10:31:51 +04:00
|
|
|
unsigned long min_unmapped_pages;
|
2006-09-26 10:31:52 +04:00
|
|
|
unsigned long min_slab_pages;
|
2005-06-22 04:14:47 +04:00
|
|
|
#endif
|
2010-02-02 08:38:57 +03:00
|
|
|
struct per_cpu_pageset __percpu *pageset;
|
2005-04-17 02:20:36 +04:00
|
|
|
/*
|
|
|
|
* free areas of different sizes
|
|
|
|
*/
|
|
|
|
spinlock_t lock;
|
2012-10-09 03:32:41 +04:00
|
|
|
#if defined CONFIG_COMPACTION || defined CONFIG_CMA
|
2012-10-09 03:32:47 +04:00
|
|
|
/* Set to true when the PG_migrate_skip bits should be cleared */
|
|
|
|
bool compact_blockskip_flush;
|
2012-10-09 03:32:45 +04:00
|
|
|
|
|
|
|
/* pfns where compaction scanners should start */
|
|
|
|
unsigned long compact_cached_free_pfn;
|
|
|
|
unsigned long compact_cached_migrate_pfn;
|
2012-10-09 03:32:41 +04:00
|
|
|
#endif
|
2005-10-30 04:16:53 +03:00
|
|
|
#ifdef CONFIG_MEMORY_HOTPLUG
|
|
|
|
/* see spanned/present_pages for more description */
|
|
|
|
seqlock_t span_seqlock;
|
|
|
|
#endif
|
2005-04-17 02:20:36 +04:00
|
|
|
struct free_area free_area[MAX_ORDER];
|
|
|
|
|
Add a bitmap that is used to track flags affecting a block of pages
Here is the latest revision of the anti-fragmentation patches. Of particular
note in this version is special treatment of high-order atomic allocations.
Care is taken to group them together and avoid grouping pages of other types
near them. Artifical tests imply that it works. I'm trying to get the
hardware together that would allow setting up of a "real" test. If anyone
already has a setup and test that can trigger the atomic-allocation problem,
I'd appreciate a test of these patches and a report. The second major change
is that these patches will apply cleanly with patches that implement
anti-fragmentation through zones.
kernbench shows effectively no performance difference varying between -0.2%
and +2% on a variety of test machines. Success rates for huge page allocation
are dramatically increased. For example, on a ppc64 machine, the vanilla
kernel was only able to allocate 1% of memory as a hugepage and this was due
to a single hugepage reserved as min_free_kbytes. With these patches applied,
17% was allocatable as superpages. With reclaim-related fixes from Andy
Whitcroft, it was 40% and further reclaim-related improvements should increase
this further.
Changelog Since V28
o Group high-order atomic allocations together
o It is no longer required to set min_free_kbytes to 10% of memory. A value
of 16384 in most cases will be sufficient
o Now applied with zone-based anti-fragmentation
o Fix incorrect VM_BUG_ON within buffered_rmqueue()
o Reorder the stack so later patches do not back out work from earlier patches
o Fix bug were journal pages were being treated as movable
o Bias placement of non-movable pages to lower PFNs
o More agressive clustering of reclaimable pages in reactions to workloads
like updatedb that flood the size of inode caches
Changelog Since V27
o Renamed anti-fragmentation to Page Clustering. Anti-fragmentation was giving
the mistaken impression that it was the 100% solution for high order
allocations. Instead, it greatly increases the chances high-order
allocations will succeed and lays the foundation for defragmentation and
memory hot-remove to work properly
o Redefine page groupings based on ability to migrate or reclaim instead of
basing on reclaimability alone
o Get rid of spurious inits
o Per-cpu lists are no longer split up per-type. Instead the per-cpu list is
searched for a page of the appropriate type
o Added more explanation commentary
o Fix up bug in pageblock code where bitmap was used before being initalised
Changelog Since V26
o Fix double init of lists in setup_pageset
Changelog Since V25
o Fix loop order of for_each_rclmtype_order so that order of loop matches args
o gfpflags_to_rclmtype uses gfp_t instead of unsigned long
o Rename get_pageblock_type() to get_page_rclmtype()
o Fix alignment problem in move_freepages()
o Add mechanism for assigning flags to blocks of pages instead of page->flags
o On fallback, do not examine the preferred list of free pages a second time
The purpose of these patches is to reduce external fragmentation by grouping
pages of related types together. When pages are migrated (or reclaimed under
memory pressure), large contiguous pages will be freed.
This patch works by categorising allocations by their ability to migrate;
Movable - The pages may be moved with the page migration mechanism. These are
generally userspace pages.
Reclaimable - These are allocations for some kernel caches that are
reclaimable or allocations that are known to be very short-lived.
Unmovable - These are pages that are allocated by the kernel that
are not trivially reclaimed. For example, the memory allocated for a
loaded module would be in this category. By default, allocations are
considered to be of this type
HighAtomic - These are high-order allocations belonging to callers that
cannot sleep or perform any IO. In practice, this is restricted to
jumbo frame allocation for network receive. It is assumed that the
allocations are short-lived
Instead of having one MAX_ORDER-sized array of free lists in struct free_area,
there is one for each type of reclaimability. Once a 2^MAX_ORDER block of
pages is split for a type of allocation, it is added to the free-lists for
that type, in effect reserving it. Hence, over time, pages of the different
types can be clustered together.
When the preferred freelists are expired, the largest possible block is taken
from an alternative list. Buddies that are split from that large block are
placed on the preferred allocation-type freelists to mitigate fragmentation.
This implementation gives best-effort for low fragmentation in all zones.
Ideally, min_free_kbytes needs to be set to a value equal to 4 * (1 <<
(MAX_ORDER-1)) pages in most cases. This would be 16384 on x86 and x86_64 for
example.
Our tests show that about 60-70% of physical memory can be allocated on a
desktop after a few days uptime. In benchmarks and stress tests, we are
finding that 80% of memory is available as contiguous blocks at the end of the
test. To compare, a standard kernel was getting < 1% of memory as large pages
on a desktop and about 8-12% of memory as large pages at the end of stress
tests.
Following this email are 12 patches that implement thie page grouping feature.
The first patch introduces a mechanism for storing flags related to a whole
block of pages. Then allocations are split between movable and all other
allocations. Following that are patches to deal with per-cpu pages and make
the mechanism configurable. The next patch moves free pages between lists
when partially allocated blocks are used for pages of another migrate type.
The second last patch groups reclaimable kernel allocations such as inode
caches together. The final patch related to groupings keeps high-order atomic
allocations.
The last two patches are more concerned with control of fragmentation. The
second last patch biases placement of non-movable allocations towards the
start of memory. This is with a view of supporting memory hot-remove of DIMMs
with higher PFNs in the future. The biasing could be enforced a lot heavier
but it would cost. The last patch agressively clusters reclaimable pages like
inode caches together.
The fragmentation reduction strategy needs to track if pages within a block
can be moved or reclaimed so that pages are freed to the appropriate list.
This patch adds a bitmap for flags affecting a whole a MAX_ORDER block of
pages.
In non-SPARSEMEM configurations, the bitmap is stored in the struct zone and
allocated during initialisation. SPARSEMEM statically allocates the bitmap in
a struct mem_section so that bitmaps do not have to be resized during memory
hotadd. This wastes a small amount of memory per unused section (usually
sizeof(unsigned long)) but the complexity of dynamically allocating the memory
is quite high.
Additional credit to Andy Whitcroft who reviewed up an earlier implementation
of the mechanism an suggested how to make it a *lot* cleaner.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Cc: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:25:47 +04:00
|
|
|
#ifndef CONFIG_SPARSEMEM
|
|
|
|
/*
|
2007-10-16 12:26:01 +04:00
|
|
|
* Flags for a pageblock_nr_pages block. See pageblock-flags.h.
|
Add a bitmap that is used to track flags affecting a block of pages
Here is the latest revision of the anti-fragmentation patches. Of particular
note in this version is special treatment of high-order atomic allocations.
Care is taken to group them together and avoid grouping pages of other types
near them. Artifical tests imply that it works. I'm trying to get the
hardware together that would allow setting up of a "real" test. If anyone
already has a setup and test that can trigger the atomic-allocation problem,
I'd appreciate a test of these patches and a report. The second major change
is that these patches will apply cleanly with patches that implement
anti-fragmentation through zones.
kernbench shows effectively no performance difference varying between -0.2%
and +2% on a variety of test machines. Success rates for huge page allocation
are dramatically increased. For example, on a ppc64 machine, the vanilla
kernel was only able to allocate 1% of memory as a hugepage and this was due
to a single hugepage reserved as min_free_kbytes. With these patches applied,
17% was allocatable as superpages. With reclaim-related fixes from Andy
Whitcroft, it was 40% and further reclaim-related improvements should increase
this further.
Changelog Since V28
o Group high-order atomic allocations together
o It is no longer required to set min_free_kbytes to 10% of memory. A value
of 16384 in most cases will be sufficient
o Now applied with zone-based anti-fragmentation
o Fix incorrect VM_BUG_ON within buffered_rmqueue()
o Reorder the stack so later patches do not back out work from earlier patches
o Fix bug were journal pages were being treated as movable
o Bias placement of non-movable pages to lower PFNs
o More agressive clustering of reclaimable pages in reactions to workloads
like updatedb that flood the size of inode caches
Changelog Since V27
o Renamed anti-fragmentation to Page Clustering. Anti-fragmentation was giving
the mistaken impression that it was the 100% solution for high order
allocations. Instead, it greatly increases the chances high-order
allocations will succeed and lays the foundation for defragmentation and
memory hot-remove to work properly
o Redefine page groupings based on ability to migrate or reclaim instead of
basing on reclaimability alone
o Get rid of spurious inits
o Per-cpu lists are no longer split up per-type. Instead the per-cpu list is
searched for a page of the appropriate type
o Added more explanation commentary
o Fix up bug in pageblock code where bitmap was used before being initalised
Changelog Since V26
o Fix double init of lists in setup_pageset
Changelog Since V25
o Fix loop order of for_each_rclmtype_order so that order of loop matches args
o gfpflags_to_rclmtype uses gfp_t instead of unsigned long
o Rename get_pageblock_type() to get_page_rclmtype()
o Fix alignment problem in move_freepages()
o Add mechanism for assigning flags to blocks of pages instead of page->flags
o On fallback, do not examine the preferred list of free pages a second time
The purpose of these patches is to reduce external fragmentation by grouping
pages of related types together. When pages are migrated (or reclaimed under
memory pressure), large contiguous pages will be freed.
This patch works by categorising allocations by their ability to migrate;
Movable - The pages may be moved with the page migration mechanism. These are
generally userspace pages.
Reclaimable - These are allocations for some kernel caches that are
reclaimable or allocations that are known to be very short-lived.
Unmovable - These are pages that are allocated by the kernel that
are not trivially reclaimed. For example, the memory allocated for a
loaded module would be in this category. By default, allocations are
considered to be of this type
HighAtomic - These are high-order allocations belonging to callers that
cannot sleep or perform any IO. In practice, this is restricted to
jumbo frame allocation for network receive. It is assumed that the
allocations are short-lived
Instead of having one MAX_ORDER-sized array of free lists in struct free_area,
there is one for each type of reclaimability. Once a 2^MAX_ORDER block of
pages is split for a type of allocation, it is added to the free-lists for
that type, in effect reserving it. Hence, over time, pages of the different
types can be clustered together.
When the preferred freelists are expired, the largest possible block is taken
from an alternative list. Buddies that are split from that large block are
placed on the preferred allocation-type freelists to mitigate fragmentation.
This implementation gives best-effort for low fragmentation in all zones.
Ideally, min_free_kbytes needs to be set to a value equal to 4 * (1 <<
(MAX_ORDER-1)) pages in most cases. This would be 16384 on x86 and x86_64 for
example.
Our tests show that about 60-70% of physical memory can be allocated on a
desktop after a few days uptime. In benchmarks and stress tests, we are
finding that 80% of memory is available as contiguous blocks at the end of the
test. To compare, a standard kernel was getting < 1% of memory as large pages
on a desktop and about 8-12% of memory as large pages at the end of stress
tests.
Following this email are 12 patches that implement thie page grouping feature.
The first patch introduces a mechanism for storing flags related to a whole
block of pages. Then allocations are split between movable and all other
allocations. Following that are patches to deal with per-cpu pages and make
the mechanism configurable. The next patch moves free pages between lists
when partially allocated blocks are used for pages of another migrate type.
The second last patch groups reclaimable kernel allocations such as inode
caches together. The final patch related to groupings keeps high-order atomic
allocations.
The last two patches are more concerned with control of fragmentation. The
second last patch biases placement of non-movable allocations towards the
start of memory. This is with a view of supporting memory hot-remove of DIMMs
with higher PFNs in the future. The biasing could be enforced a lot heavier
but it would cost. The last patch agressively clusters reclaimable pages like
inode caches together.
The fragmentation reduction strategy needs to track if pages within a block
can be moved or reclaimed so that pages are freed to the appropriate list.
This patch adds a bitmap for flags affecting a whole a MAX_ORDER block of
pages.
In non-SPARSEMEM configurations, the bitmap is stored in the struct zone and
allocated during initialisation. SPARSEMEM statically allocates the bitmap in
a struct mem_section so that bitmaps do not have to be resized during memory
hotadd. This wastes a small amount of memory per unused section (usually
sizeof(unsigned long)) but the complexity of dynamically allocating the memory
is quite high.
Additional credit to Andy Whitcroft who reviewed up an earlier implementation
of the mechanism an suggested how to make it a *lot* cleaner.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Cc: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:25:47 +04:00
|
|
|
* In SPARSEMEM, this map is stored in struct mem_section
|
|
|
|
*/
|
|
|
|
unsigned long *pageblock_flags;
|
|
|
|
#endif /* CONFIG_SPARSEMEM */
|
|
|
|
|
2010-05-25 01:32:32 +04:00
|
|
|
#ifdef CONFIG_COMPACTION
|
|
|
|
/*
|
|
|
|
* On compaction failure, 1<<compact_defer_shift compactions
|
|
|
|
* are skipped before trying again. The number attempted since
|
|
|
|
* last failure is tracked with compact_considered.
|
|
|
|
*/
|
|
|
|
unsigned int compact_considered;
|
|
|
|
unsigned int compact_defer_shift;
|
2012-03-22 03:33:52 +04:00
|
|
|
int compact_order_failed;
|
2010-05-25 01:32:32 +04:00
|
|
|
#endif
|
2005-04-17 02:20:36 +04:00
|
|
|
|
|
|
|
ZONE_PADDING(_pad1_)
|
|
|
|
|
|
|
|
/* Fields commonly accessed by the page reclaim scanner */
|
2012-01-13 05:18:10 +04:00
|
|
|
spinlock_t lru_lock;
|
|
|
|
struct lruvec lruvec;
|
2008-10-19 07:26:32 +04:00
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
unsigned long pages_scanned; /* since last reclaim */
|
2007-10-17 10:25:54 +04:00
|
|
|
unsigned long flags; /* zone flags, see below */
|
2005-06-22 04:14:41 +04:00
|
|
|
|
2006-06-30 12:55:33 +04:00
|
|
|
/* Zone statistics */
|
|
|
|
atomic_long_t vm_stat[NR_VM_ZONE_STAT_ITEMS];
|
2006-01-19 04:42:31 +03:00
|
|
|
|
2008-10-19 07:26:34 +04:00
|
|
|
/*
|
|
|
|
* The target ratio of ACTIVE_ANON to INACTIVE_ANON pages on
|
|
|
|
* this zone's LRU. Maintained by the pageout code.
|
|
|
|
*/
|
|
|
|
unsigned int inactive_ratio;
|
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
|
|
|
|
ZONE_PADDING(_pad2_)
|
|
|
|
/* Rarely used or read-mostly fields */
|
|
|
|
|
|
|
|
/*
|
|
|
|
* wait_table -- the array holding the hash table
|
2006-06-23 13:03:08 +04:00
|
|
|
* wait_table_hash_nr_entries -- the size of the hash table array
|
2005-04-17 02:20:36 +04:00
|
|
|
* wait_table_bits -- wait_table_size == (1 << wait_table_bits)
|
|
|
|
*
|
|
|
|
* The purpose of all these is to keep track of the people
|
|
|
|
* waiting for a page to become available and make them
|
|
|
|
* runnable again when possible. The trouble is that this
|
|
|
|
* consumes a lot of space, especially when so few things
|
|
|
|
* wait on pages at a given time. So instead of using
|
|
|
|
* per-page waitqueues, we use a waitqueue hash table.
|
|
|
|
*
|
|
|
|
* The bucket discipline is to sleep on the same queue when
|
|
|
|
* colliding and wake all in that wait queue when removing.
|
|
|
|
* When something wakes, it must check to be sure its page is
|
|
|
|
* truly available, a la thundering herd. The cost of a
|
|
|
|
* collision is great, but given the expected load of the
|
|
|
|
* table, they should be so rare as to be outweighed by the
|
|
|
|
* benefits from the saved space.
|
|
|
|
*
|
|
|
|
* __wait_on_page_locked() and unlock_page() in mm/filemap.c, are the
|
|
|
|
* primary users of these fields, and in mm/page_alloc.c
|
|
|
|
* free_area_init_core() performs the initialization of them.
|
|
|
|
*/
|
|
|
|
wait_queue_head_t * wait_table;
|
2006-06-23 13:03:08 +04:00
|
|
|
unsigned long wait_table_hash_nr_entries;
|
2005-04-17 02:20:36 +04:00
|
|
|
unsigned long wait_table_bits;
|
|
|
|
|
|
|
|
/*
|
|
|
|
* Discontig memory support fields.
|
|
|
|
*/
|
|
|
|
struct pglist_data *zone_pgdat;
|
|
|
|
/* zone_start_pfn == zone_start_paddr >> PAGE_SHIFT */
|
|
|
|
unsigned long zone_start_pfn;
|
|
|
|
|
2005-10-30 04:16:53 +03:00
|
|
|
/*
|
2012-12-13 01:52:12 +04:00
|
|
|
* spanned_pages is the total pages spanned by the zone, including
|
|
|
|
* holes, which is calculated as:
|
|
|
|
* spanned_pages = zone_end_pfn - zone_start_pfn;
|
2005-10-30 04:16:53 +03:00
|
|
|
*
|
2012-12-13 01:52:12 +04:00
|
|
|
* present_pages is physical pages existing within the zone, which
|
|
|
|
* is calculated as:
|
2013-03-26 21:30:44 +04:00
|
|
|
* present_pages = spanned_pages - absent_pages(pages in holes);
|
2012-12-13 01:52:12 +04:00
|
|
|
*
|
|
|
|
* managed_pages is present pages managed by the buddy system, which
|
|
|
|
* is calculated as (reserved_pages includes pages allocated by the
|
|
|
|
* bootmem allocator):
|
|
|
|
* managed_pages = present_pages - reserved_pages;
|
|
|
|
*
|
|
|
|
* So present_pages may be used by memory hotplug or memory power
|
|
|
|
* management logic to figure out unmanaged pages by checking
|
|
|
|
* (present_pages - managed_pages). And managed_pages should be used
|
|
|
|
* by page allocator and vm scanner to calculate all kinds of watermarks
|
|
|
|
* and thresholds.
|
|
|
|
*
|
|
|
|
* Locking rules:
|
|
|
|
*
|
|
|
|
* zone_start_pfn and spanned_pages are protected by span_seqlock.
|
|
|
|
* It is a seqlock because it has to be read outside of zone->lock,
|
|
|
|
* and it is done in the main allocator path. But, it is written
|
|
|
|
* quite infrequently.
|
|
|
|
*
|
|
|
|
* The span_seq lock is declared along with zone->lock because it is
|
2005-10-30 04:16:53 +03:00
|
|
|
* frequently read in proximity to zone->lock. It's good to
|
|
|
|
* give them a chance of being in the same cacheline.
|
2012-12-13 01:52:12 +04:00
|
|
|
*
|
2013-07-04 02:03:14 +04:00
|
|
|
* Write access to present_pages at runtime should be protected by
|
|
|
|
* lock_memory_hotplug()/unlock_memory_hotplug(). Any reader who can't
|
|
|
|
* tolerant drift of present_pages should hold memory hotplug lock to
|
|
|
|
* get a stable value.
|
|
|
|
*
|
|
|
|
* Read access to managed_pages should be safe because it's unsigned
|
|
|
|
* long. Write access to zone->managed_pages and totalram_pages are
|
|
|
|
* protected by managed_page_count_lock at runtime. Idealy only
|
|
|
|
* adjust_managed_page_count() should be used instead of directly
|
|
|
|
* touching zone->managed_pages and totalram_pages.
|
2005-10-30 04:16:53 +03:00
|
|
|
*/
|
2012-12-13 01:52:12 +04:00
|
|
|
unsigned long spanned_pages;
|
|
|
|
unsigned long present_pages;
|
|
|
|
unsigned long managed_pages;
|
2005-04-17 02:20:36 +04:00
|
|
|
|
|
|
|
/*
|
|
|
|
* rarely used fields:
|
|
|
|
*/
|
2006-12-07 07:40:36 +03:00
|
|
|
const char *name;
|
2006-01-08 12:01:27 +03:00
|
|
|
} ____cacheline_internodealigned_in_smp;
|
2005-04-17 02:20:36 +04:00
|
|
|
|
2007-10-17 10:25:54 +04:00
|
|
|
typedef enum {
|
|
|
|
ZONE_RECLAIM_LOCKED, /* prevents concurrent reclaim */
|
2007-10-17 10:25:55 +04:00
|
|
|
ZONE_OOM_LOCKED, /* zone is in OOM killer zonelist */
|
2010-10-27 01:21:45 +04:00
|
|
|
ZONE_CONGESTED, /* zone has many dirty pages backed by
|
|
|
|
* a congested BDI
|
|
|
|
*/
|
2013-07-04 02:01:50 +04:00
|
|
|
ZONE_TAIL_LRU_DIRTY, /* reclaim scanning has recently found
|
|
|
|
* many dirty file pages at the tail
|
|
|
|
* of the LRU.
|
|
|
|
*/
|
2013-07-04 02:01:51 +04:00
|
|
|
ZONE_WRITEBACK, /* reclaim scanning has recently found
|
|
|
|
* many pages under writeback
|
|
|
|
*/
|
2007-10-17 10:25:54 +04:00
|
|
|
} zone_flags_t;
|
|
|
|
|
|
|
|
static inline void zone_set_flag(struct zone *zone, zone_flags_t flag)
|
|
|
|
{
|
|
|
|
set_bit(flag, &zone->flags);
|
|
|
|
}
|
2007-10-17 10:26:01 +04:00
|
|
|
|
|
|
|
static inline int zone_test_and_set_flag(struct zone *zone, zone_flags_t flag)
|
|
|
|
{
|
|
|
|
return test_and_set_bit(flag, &zone->flags);
|
|
|
|
}
|
|
|
|
|
2007-10-17 10:25:54 +04:00
|
|
|
static inline void zone_clear_flag(struct zone *zone, zone_flags_t flag)
|
|
|
|
{
|
|
|
|
clear_bit(flag, &zone->flags);
|
|
|
|
}
|
|
|
|
|
2010-10-27 01:21:45 +04:00
|
|
|
static inline int zone_is_reclaim_congested(const struct zone *zone)
|
|
|
|
{
|
|
|
|
return test_bit(ZONE_CONGESTED, &zone->flags);
|
|
|
|
}
|
|
|
|
|
2013-07-04 02:01:50 +04:00
|
|
|
static inline int zone_is_reclaim_dirty(const struct zone *zone)
|
|
|
|
{
|
|
|
|
return test_bit(ZONE_TAIL_LRU_DIRTY, &zone->flags);
|
|
|
|
}
|
|
|
|
|
2013-07-04 02:01:51 +04:00
|
|
|
static inline int zone_is_reclaim_writeback(const struct zone *zone)
|
|
|
|
{
|
|
|
|
return test_bit(ZONE_WRITEBACK, &zone->flags);
|
|
|
|
}
|
|
|
|
|
2007-10-17 10:25:54 +04:00
|
|
|
static inline int zone_is_reclaim_locked(const struct zone *zone)
|
|
|
|
{
|
|
|
|
return test_bit(ZONE_RECLAIM_LOCKED, &zone->flags);
|
|
|
|
}
|
2007-10-17 10:26:01 +04:00
|
|
|
|
2007-10-17 10:25:55 +04:00
|
|
|
static inline int zone_is_oom_locked(const struct zone *zone)
|
|
|
|
{
|
|
|
|
return test_bit(ZONE_OOM_LOCKED, &zone->flags);
|
|
|
|
}
|
2007-10-17 10:25:54 +04:00
|
|
|
|
2013-03-23 02:04:43 +04:00
|
|
|
static inline unsigned long zone_end_pfn(const struct zone *zone)
|
2013-02-23 04:35:23 +04:00
|
|
|
{
|
|
|
|
return zone->zone_start_pfn + zone->spanned_pages;
|
|
|
|
}
|
|
|
|
|
|
|
|
static inline bool zone_spans_pfn(const struct zone *zone, unsigned long pfn)
|
|
|
|
{
|
|
|
|
return zone->zone_start_pfn <= pfn && pfn < zone_end_pfn(zone);
|
|
|
|
}
|
|
|
|
|
2013-02-23 04:35:24 +04:00
|
|
|
static inline bool zone_is_initialized(struct zone *zone)
|
|
|
|
{
|
|
|
|
return !!zone->wait_table;
|
|
|
|
}
|
|
|
|
|
|
|
|
static inline bool zone_is_empty(struct zone *zone)
|
|
|
|
{
|
|
|
|
return zone->spanned_pages == 0;
|
|
|
|
}
|
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
/*
|
|
|
|
* The "priority" of VM scanning is how much of the queues we will scan in one
|
|
|
|
* go. A value of 12 for DEF_PRIORITY implies that we will scan 1/4096th of the
|
|
|
|
* queues ("queue_length >> 12") during an aging round.
|
|
|
|
*/
|
|
|
|
#define DEF_PRIORITY 12
|
|
|
|
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
/* Maximum number of zones on a zonelist */
|
|
|
|
#define MAX_ZONES_PER_ZONELIST (MAX_NUMNODES * MAX_NR_ZONES)
|
|
|
|
|
|
|
|
#ifdef CONFIG_NUMA
|
2007-10-16 12:25:37 +04:00
|
|
|
|
|
|
|
/*
|
2011-02-04 08:43:48 +03:00
|
|
|
* The NUMA zonelists are doubled because we need zonelists that restrict the
|
2007-10-16 12:25:37 +04:00
|
|
|
* allocations to a single node for GFP_THISNODE.
|
|
|
|
*
|
2008-04-28 13:12:16 +04:00
|
|
|
* [0] : Zonelist with fallback
|
|
|
|
* [1] : No fallback (GFP_THISNODE)
|
2007-10-16 12:25:37 +04:00
|
|
|
*/
|
2008-04-28 13:12:16 +04:00
|
|
|
#define MAX_ZONELISTS 2
|
2007-10-16 12:25:37 +04:00
|
|
|
|
|
|
|
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
/*
|
|
|
|
* We cache key information from each zonelist for smaller cache
|
|
|
|
* footprint when scanning for free pages in get_page_from_freelist().
|
|
|
|
*
|
|
|
|
* 1) The BITMAP fullzones tracks which zones in a zonelist have come
|
|
|
|
* up short of free memory since the last time (last_fullzone_zap)
|
|
|
|
* we zero'd fullzones.
|
|
|
|
* 2) The array z_to_n[] maps each zone in the zonelist to its node
|
|
|
|
* id, so that we can efficiently evaluate whether that node is
|
|
|
|
* set in the current tasks mems_allowed.
|
|
|
|
*
|
|
|
|
* Both fullzones and z_to_n[] are one-to-one with the zonelist,
|
|
|
|
* indexed by a zones offset in the zonelist zones[] array.
|
|
|
|
*
|
|
|
|
* The get_page_from_freelist() routine does two scans. During the
|
|
|
|
* first scan, we skip zones whose corresponding bit in 'fullzones'
|
|
|
|
* is set or whose corresponding node in current->mems_allowed (which
|
|
|
|
* comes from cpusets) is not set. During the second scan, we bypass
|
|
|
|
* this zonelist_cache, to ensure we look methodically at each zone.
|
|
|
|
*
|
|
|
|
* Once per second, we zero out (zap) fullzones, forcing us to
|
|
|
|
* reconsider nodes that might have regained more free memory.
|
|
|
|
* The field last_full_zap is the time we last zapped fullzones.
|
|
|
|
*
|
|
|
|
* This mechanism reduces the amount of time we waste repeatedly
|
|
|
|
* reexaming zones for free memory when they just came up low on
|
|
|
|
* memory momentarilly ago.
|
|
|
|
*
|
|
|
|
* The zonelist_cache struct members logically belong in struct
|
|
|
|
* zonelist. However, the mempolicy zonelists constructed for
|
|
|
|
* MPOL_BIND are intentionally variable length (and usually much
|
|
|
|
* shorter). A general purpose mechanism for handling structs with
|
|
|
|
* multiple variable length members is more mechanism than we want
|
|
|
|
* here. We resort to some special case hackery instead.
|
|
|
|
*
|
|
|
|
* The MPOL_BIND zonelists don't need this zonelist_cache (in good
|
|
|
|
* part because they are shorter), so we put the fixed length stuff
|
|
|
|
* at the front of the zonelist struct, ending in a variable length
|
|
|
|
* zones[], as is needed by MPOL_BIND.
|
|
|
|
*
|
|
|
|
* Then we put the optional zonelist cache on the end of the zonelist
|
|
|
|
* struct. This optional stuff is found by a 'zlcache_ptr' pointer in
|
|
|
|
* the fixed length portion at the front of the struct. This pointer
|
|
|
|
* both enables us to find the zonelist cache, and in the case of
|
|
|
|
* MPOL_BIND zonelists, (which will just set the zlcache_ptr to NULL)
|
|
|
|
* to know that the zonelist cache is not there.
|
|
|
|
*
|
|
|
|
* The end result is that struct zonelists come in two flavors:
|
|
|
|
* 1) The full, fixed length version, shown below, and
|
|
|
|
* 2) The custom zonelists for MPOL_BIND.
|
|
|
|
* The custom MPOL_BIND zonelists have a NULL zlcache_ptr and no zlcache.
|
|
|
|
*
|
|
|
|
* Even though there may be multiple CPU cores on a node modifying
|
|
|
|
* fullzones or last_full_zap in the same zonelist_cache at the same
|
|
|
|
* time, we don't lock it. This is just hint data - if it is wrong now
|
|
|
|
* and then, the allocator will still function, perhaps a bit slower.
|
|
|
|
*/
|
|
|
|
|
|
|
|
|
|
|
|
struct zonelist_cache {
|
|
|
|
unsigned short z_to_n[MAX_ZONES_PER_ZONELIST]; /* zone->nid */
|
2006-12-07 07:31:49 +03:00
|
|
|
DECLARE_BITMAP(fullzones, MAX_ZONES_PER_ZONELIST); /* zone full? */
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
unsigned long last_full_zap; /* when last zap'd (jiffies) */
|
|
|
|
};
|
|
|
|
#else
|
2008-04-28 13:12:16 +04:00
|
|
|
#define MAX_ZONELISTS 1
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
struct zonelist_cache;
|
|
|
|
#endif
|
|
|
|
|
2008-04-28 13:12:17 +04:00
|
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|
/*
|
|
|
|
* This struct contains information about a zone in a zonelist. It is stored
|
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|
|
* here to avoid dereferences into large structures and lookups of tables
|
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|
|
*/
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|
|
struct zoneref {
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struct zone *zone; /* Pointer to actual zone */
|
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int zone_idx; /* zone_idx(zoneref->zone) */
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|
};
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|
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|
2005-04-17 02:20:36 +04:00
|
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|
/*
|
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|
|
* One allocation request operates on a zonelist. A zonelist
|
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|
* is a list of zones, the first one is the 'goal' of the
|
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|
|
* allocation, the other zones are fallback zones, in decreasing
|
|
|
|
* priority.
|
|
|
|
*
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
* If zlcache_ptr is not NULL, then it is just the address of zlcache,
|
|
|
|
* as explained above. If zlcache_ptr is NULL, there is no zlcache.
|
2008-04-28 13:12:17 +04:00
|
|
|
* *
|
|
|
|
* To speed the reading of the zonelist, the zonerefs contain the zone index
|
|
|
|
* of the entry being read. Helper functions to access information given
|
|
|
|
* a struct zoneref are
|
|
|
|
*
|
|
|
|
* zonelist_zone() - Return the struct zone * for an entry in _zonerefs
|
|
|
|
* zonelist_zone_idx() - Return the index of the zone for an entry
|
|
|
|
* zonelist_node_idx() - Return the index of the node for an entry
|
2005-04-17 02:20:36 +04:00
|
|
|
*/
|
|
|
|
struct zonelist {
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
struct zonelist_cache *zlcache_ptr; // NULL or &zlcache
|
2008-04-28 13:12:17 +04:00
|
|
|
struct zoneref _zonerefs[MAX_ZONES_PER_ZONELIST + 1];
|
[PATCH] memory page_alloc zonelist caching speedup
Optimize the critical zonelist scanning for free pages in the kernel memory
allocator by caching the zones that were found to be full recently, and
skipping them.
Remembers the zones in a zonelist that were short of free memory in the
last second. And it stashes a zone-to-node table in the zonelist struct,
to optimize that conversion (minimize its cache footprint.)
Recent changes:
This differs in a significant way from a similar patch that I
posted a week ago. Now, instead of having a nodemask_t of
recently full nodes, I have a bitmask of recently full zones.
This solves a problem that last weeks patch had, which on
systems with multiple zones per node (such as DMA zone) would
take seeing any of these zones full as meaning that all zones
on that node were full.
Also I changed names - from "zonelist faster" to "zonelist cache",
as that seemed to better convey what we're doing here - caching
some of the key zonelist state (for faster access.)
See below for some performance benchmark results. After all that
discussion with David on why I didn't need them, I went and got
some ;). I wanted to verify that I had not hurt the normal case
of memory allocation noticeably. At least for my one little
microbenchmark, I found (1) the normal case wasn't affected, and
(2) workloads that forced scanning across multiple nodes for
memory improved up to 10% fewer System CPU cycles and lower
elapsed clock time ('sys' and 'real'). Good. See details, below.
I didn't have the logic in get_page_from_freelist() for various
full nodes and zone reclaim failures correct. That should be
fixed up now - notice the new goto labels zonelist_scan,
this_zone_full, and try_next_zone, in get_page_from_freelist().
There are two reasons I persued this alternative, over some earlier
proposals that would have focused on optimizing the fake numa
emulation case by caching the last useful zone:
1) Contrary to what I said before, we (SGI, on large ia64 sn2 systems)
have seen real customer loads where the cost to scan the zonelist
was a problem, due to many nodes being full of memory before
we got to a node we could use. Or at least, I think we have.
This was related to me by another engineer, based on experiences
from some time past. So this is not guaranteed. Most likely, though.
The following approach should help such real numa systems just as
much as it helps fake numa systems, or any combination thereof.
2) The effort to distinguish fake from real numa, using node_distance,
so that we could cache a fake numa node and optimize choosing
it over equivalent distance fake nodes, while continuing to
properly scan all real nodes in distance order, was going to
require a nasty blob of zonelist and node distance munging.
The following approach has no new dependency on node distances or
zone sorting.
See comment in the patch below for a description of what it actually does.
Technical details of note (or controversy):
- See the use of "zlc_active" and "did_zlc_setup" below, to delay
adding any work for this new mechanism until we've looked at the
first zone in zonelist. I figured the odds of the first zone
having the memory we needed were high enough that we should just
look there, first, then get fancy only if we need to keep looking.
- Some odd hackery was needed to add items to struct zonelist, while
not tripping up the custom zonelists built by the mm/mempolicy.c
code for MPOL_BIND. My usual wordy comments below explain this.
Search for "MPOL_BIND".
- Some per-node data in the struct zonelist is now modified frequently,
with no locking. Multiple CPU cores on a node could hit and mangle
this data. The theory is that this is just performance hint data,
and the memory allocator will work just fine despite any such mangling.
The fields at risk are the struct 'zonelist_cache' fields 'fullzones'
(a bitmask) and 'last_full_zap' (unsigned long jiffies). It should
all be self correcting after at most a one second delay.
- This still does a linear scan of the same lengths as before. All
I've optimized is making the scan faster, not algorithmically
shorter. It is now able to scan a compact array of 'unsigned
short' in the case of many full nodes, so one cache line should
cover quite a few nodes, rather than each node hitting another
one or two new and distinct cache lines.
- If both Andi and Nick don't find this too complicated, I will be
(pleasantly) flabbergasted.
- I removed the comment claiming we only use one cachline's worth of
zonelist. We seem, at least in the fake numa case, to have put the
lie to that claim.
- I pay no attention to the various watermarks and such in this performance
hint. A node could be marked full for one watermark, and then skipped
over when searching for a page using a different watermark. I think
that's actually quite ok, as it will tend to slightly increase the
spreading of memory over other nodes, away from a memory stressed node.
===============
Performance - some benchmark results and analysis:
This benchmark runs a memory hog program that uses multiple
threads to touch alot of memory as quickly as it can.
Multiple runs were made, touching 12, 38, 64 or 90 GBytes out of
the total 96 GBytes on the system, and using 1, 19, 37, or 55
threads (on a 56 CPU system.) System, user and real (elapsed)
timings were recorded for each run, shown in units of seconds,
in the table below.
Two kernels were tested - 2.6.18-mm3 and the same kernel with
this zonelist caching patch added. The table also shows the
percentage improvement the zonelist caching sys time is over
(lower than) the stock *-mm kernel.
number 2.6.18-mm3 zonelist-cache delta (< 0 good) percent
GBs N ------------ -------------- ---------------- systime
mem threads sys user real sys user real sys user real better
12 1 153 24 177 151 24 176 -2 0 -1 1%
12 19 99 22 8 99 22 8 0 0 0 0%
12 37 111 25 6 112 25 6 1 0 0 -0%
12 55 115 25 5 110 23 5 -5 -2 0 4%
38 1 502 74 576 497 73 570 -5 -1 -6 0%
38 19 426 78 48 373 76 39 -53 -2 -9 12%
38 37 544 83 36 547 82 36 3 -1 0 -0%
38 55 501 77 23 511 80 24 10 3 1 -1%
64 1 917 125 1042 890 124 1014 -27 -1 -28 2%
64 19 1118 138 119 965 141 103 -153 3 -16 13%
64 37 1202 151 94 1136 150 81 -66 -1 -13 5%
64 55 1118 141 61 1072 140 58 -46 -1 -3 4%
90 1 1342 177 1519 1275 174 1450 -67 -3 -69 4%
90 19 2392 199 192 2116 189 176 -276 -10 -16 11%
90 37 3313 238 175 2972 225 145 -341 -13 -30 10%
90 55 1948 210 104 1843 213 100 -105 3 -4 5%
Notes:
1) This test ran a memory hog program that started a specified number N of
threads, and had each thread allocate and touch 1/N'th of
the total memory to be used in the test run in a single loop,
writing a constant word to memory, one store every 4096 bytes.
Watching this test during some earlier trial runs, I would see
each of these threads sit down on one CPU and stay there, for
the remainder of the pass, a different CPU for each thread.
2) The 'real' column is not comparable to the 'sys' or 'user' columns.
The 'real' column is seconds wall clock time elapsed, from beginning
to end of that test pass. The 'sys' and 'user' columns are total
CPU seconds spent on that test pass. For a 19 thread test run,
for example, the sum of 'sys' and 'user' could be up to 19 times the
number of 'real' elapsed wall clock seconds.
3) Tests were run on a fresh, single-user boot, to minimize the amount
of memory already in use at the start of the test, and to minimize
the amount of background activity that might interfere.
4) Tests were done on a 56 CPU, 28 Node system with 96 GBytes of RAM.
5) Notice that the 'real' time gets large for the single thread runs, even
though the measured 'sys' and 'user' times are modest. I'm not sure what
that means - probably something to do with it being slow for one thread to
be accessing memory along ways away. Perhaps the fake numa system, running
ostensibly the same workload, would not show this substantial degradation
of 'real' time for one thread on many nodes -- lets hope not.
6) The high thread count passes (one thread per CPU - on 55 of 56 CPUs)
ran quite efficiently, as one might expect. Each pair of threads needed
to allocate and touch the memory on the node the two threads shared, a
pleasantly parallizable workload.
7) The intermediate thread count passes, when asking for alot of memory forcing
them to go to a few neighboring nodes, improved the most with this zonelist
caching patch.
Conclusions:
* This zonelist cache patch probably makes little difference one way or the
other for most workloads on real numa hardware, if those workloads avoid
heavy off node allocations.
* For memory intensive workloads requiring substantial off-node allocations
on real numa hardware, this patch improves both kernel and elapsed timings
up to ten per-cent.
* For fake numa systems, I'm optimistic, but will have to leave that up to
Rohit Seth to actually test (once I get him a 2.6.18 backport.)
Signed-off-by: Paul Jackson <pj@sgi.com>
Cc: Rohit Seth <rohitseth@google.com>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Cc: David Rientjes <rientjes@cs.washington.edu>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-12-07 07:31:48 +03:00
|
|
|
#ifdef CONFIG_NUMA
|
|
|
|
struct zonelist_cache zlcache; // optional ...
|
|
|
|
#endif
|
2005-04-17 02:20:36 +04:00
|
|
|
};
|
|
|
|
|
2011-12-08 22:22:09 +04:00
|
|
|
#ifdef CONFIG_HAVE_MEMBLOCK_NODE_MAP
|
[PATCH] Introduce mechanism for registering active regions of memory
At a basic level, architectures define structures to record where active
ranges of page frames are located. Once located, the code to calculate zone
sizes and holes in each architecture is very similar. Some of this zone and
hole sizing code is difficult to read for no good reason. This set of patches
eliminates the similar-looking architecture-specific code.
The patches introduce a mechanism where architectures register where the
active ranges of page frames are with add_active_range(). When all areas have
been discovered, free_area_init_nodes() is called to initialise the pgdat and
zones. The zone sizes and holes are then calculated in an architecture
independent manner.
Patch 1 introduces the mechanism for registering and initialising PFN ranges
Patch 2 changes ppc to use the mechanism - 139 arch-specific LOC removed
Patch 3 changes x86 to use the mechanism - 136 arch-specific LOC removed
Patch 4 changes x86_64 to use the mechanism - 74 arch-specific LOC removed
Patch 5 changes ia64 to use the mechanism - 52 arch-specific LOC removed
Patch 6 accounts for mem_map as a memory hole as the pages are not reclaimable.
It adjusts the watermarks slightly
Tony Luck has successfully tested for ia64 on Itanium with tiger_defconfig,
gensparse_defconfig and defconfig. Bob Picco has also tested and debugged on
IA64. Jack Steiner successfully boot tested on a mammoth SGI IA64-based
machine. These were on patches against 2.6.17-rc1 and release 3 of these
patches but there have been no ia64-changes since release 3.
There are differences in the zone sizes for x86_64 as the arch-specific code
for x86_64 accounts the kernel image and the starting mem_maps as memory holes
but the architecture-independent code accounts the memory as present.
The big benefit of this set of patches is a sizable reduction of
architecture-specific code, some of which is very hairy. There should be a
greater reduction when other architectures use the same mechanisms for zone
and hole sizing but I lack the hardware to test on.
Additional credit;
Dave Hansen for the initial suggestion and comments on early patches
Andy Whitcroft for reviewing early versions and catching numerous
errors
Tony Luck for testing and debugging on IA64
Bob Picco for fixing bugs related to pfn registration, reviewing a
number of patch revisions, providing a number of suggestions
on future direction and testing heavily
Jack Steiner and Robin Holt for testing on IA64 and clarifying
issues related to memory holes
Yasunori for testing on IA64
Andi Kleen for reviewing and feeding back about x86_64
Christian Kujau for providing valuable information related to ACPI
problems on x86_64 and testing potential fixes
This patch:
Define the structure to represent an active range of page frames within a node
in an architecture independent manner. Architectures are expected to register
active ranges of PFNs using add_active_range(nid, start_pfn, end_pfn) and call
free_area_init_nodes() passing the PFNs of the end of each zone.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Andy Whitcroft <apw@shadowen.org>
Cc: Andi Kleen <ak@muc.de>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Paul Mackerras <paulus@samba.org>
Cc: "Keith Mannthey" <kmannth@gmail.com>
Cc: "Luck, Tony" <tony.luck@intel.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-09-27 12:49:43 +04:00
|
|
|
struct node_active_region {
|
|
|
|
unsigned long start_pfn;
|
|
|
|
unsigned long end_pfn;
|
|
|
|
int nid;
|
|
|
|
};
|
2011-12-08 22:22:09 +04:00
|
|
|
#endif /* CONFIG_HAVE_MEMBLOCK_NODE_MAP */
|
2005-04-17 02:20:36 +04:00
|
|
|
|
2006-09-27 12:50:01 +04:00
|
|
|
#ifndef CONFIG_DISCONTIGMEM
|
|
|
|
/* The array of struct pages - for discontigmem use pgdat->lmem_map */
|
|
|
|
extern struct page *mem_map;
|
|
|
|
#endif
|
|
|
|
|
2005-04-17 02:20:36 +04:00
|
|
|
/*
|
|
|
|
* The pg_data_t structure is used in machines with CONFIG_DISCONTIGMEM
|
|
|
|
* (mostly NUMA machines?) to denote a higher-level memory zone than the
|
|
|
|
* zone denotes.
|
|
|
|
*
|
|
|
|
* On NUMA machines, each NUMA node would have a pg_data_t to describe
|
|
|
|
* it's memory layout.
|
|
|
|
*
|
|
|
|
* Memory statistics and page replacement data structures are maintained on a
|
|
|
|
* per-zone basis.
|
|
|
|
*/
|
|
|
|
struct bootmem_data;
|
|
|
|
typedef struct pglist_data {
|
|
|
|
struct zone node_zones[MAX_NR_ZONES];
|
2007-10-16 12:25:37 +04:00
|
|
|
struct zonelist node_zonelists[MAX_ZONELISTS];
|
2005-04-17 02:20:36 +04:00
|
|
|
int nr_zones;
|
2008-10-19 07:28:16 +04:00
|
|
|
#ifdef CONFIG_FLAT_NODE_MEM_MAP /* means !SPARSEMEM */
|
2005-04-17 02:20:36 +04:00
|
|
|
struct page *node_mem_map;
|
2012-08-01 03:43:02 +04:00
|
|
|
#ifdef CONFIG_MEMCG
|
2008-10-19 07:28:16 +04:00
|
|
|
struct page_cgroup *node_page_cgroup;
|
|
|
|
#endif
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#endif
|
2010-02-10 12:20:20 +03:00
|
|
|
#ifndef CONFIG_NO_BOOTMEM
|
2005-04-17 02:20:36 +04:00
|
|
|
struct bootmem_data *bdata;
|
2010-02-10 12:20:20 +03:00
|
|
|
#endif
|
2005-10-30 04:16:52 +03:00
|
|
|
#ifdef CONFIG_MEMORY_HOTPLUG
|
|
|
|
/*
|
|
|
|
* Must be held any time you expect node_start_pfn, node_present_pages
|
|
|
|
* or node_spanned_pages stay constant. Holding this will also
|
|
|
|
* guarantee that any pfn_valid() stays that way.
|
|
|
|
*
|
2013-07-04 02:02:09 +04:00
|
|
|
* pgdat_resize_lock() and pgdat_resize_unlock() are provided to
|
|
|
|
* manipulate node_size_lock without checking for CONFIG_MEMORY_HOTPLUG.
|
|
|
|
*
|
2013-07-04 02:02:08 +04:00
|
|
|
* Nests above zone->lock and zone->span_seqlock
|
2005-10-30 04:16:52 +03:00
|
|
|
*/
|
|
|
|
spinlock_t node_size_lock;
|
|
|
|
#endif
|
2005-04-17 02:20:36 +04:00
|
|
|
unsigned long node_start_pfn;
|
|
|
|
unsigned long node_present_pages; /* total number of physical pages */
|
|
|
|
unsigned long node_spanned_pages; /* total size of physical page
|
|
|
|
range, including holes */
|
|
|
|
int node_id;
|
2012-10-09 03:33:24 +04:00
|
|
|
nodemask_t reclaim_nodes; /* Nodes allowed to reclaim from */
|
2005-04-17 02:20:36 +04:00
|
|
|
wait_queue_head_t kswapd_wait;
|
2012-08-01 03:44:35 +04:00
|
|
|
wait_queue_head_t pfmemalloc_wait;
|
2012-07-12 01:01:52 +04:00
|
|
|
struct task_struct *kswapd; /* Protected by lock_memory_hotplug() */
|
2005-04-17 02:20:36 +04:00
|
|
|
int kswapd_max_order;
|
mm: kswapd: stop high-order balancing when any suitable zone is balanced
Simon Kirby reported the following problem
We're seeing cases on a number of servers where cache never fully
grows to use all available memory. Sometimes we see servers with 4 GB
of memory that never seem to have less than 1.5 GB free, even with a
constantly-active VM. In some cases, these servers also swap out while
this happens, even though they are constantly reading the working set
into memory. We have been seeing this happening for a long time; I
don't think it's anything recent, and it still happens on 2.6.36.
After some debugging work by Simon, Dave Hansen and others, the prevaling
theory became that kswapd is reclaiming order-3 pages requested by SLUB
too aggressive about it.
There are two apparent problems here. On the target machine, there is a
small Normal zone in comparison to DMA32. As kswapd tries to balance all
zones, it would continually try reclaiming for Normal even though DMA32
was balanced enough for callers. The second problem is that
sleeping_prematurely() does not use the same logic as balance_pgdat() when
deciding whether to sleep or not. This keeps kswapd artifically awake.
A number of tests were run and the figures from previous postings will
look very different for a few reasons. One, the old figures were forcing
my network card to use GFP_ATOMIC in attempt to replicate Simon's problem.
Second, I previous specified slub_min_order=3 again in an attempt to
reproduce Simon's problem. In this posting, I'm depending on Simon to say
whether his problem is fixed or not and these figures are to show the
impact to the ordinary cases. Finally, the "vmscan" figures are taken
from /proc/vmstat instead of the tracepoints. There is less information
but recording is less disruptive.
The first test of relevance was postmark with a process running in the
background reading a large amount of anonymous memory in blocks. The
objective was to vaguely simulate what was happening on Simon's machine
and it's memory intensive enough to have kswapd awake.
POSTMARK
traceonly kanyzone
Transactions per second: 156.00 ( 0.00%) 153.00 (-1.96%)
Data megabytes read per second: 21.51 ( 0.00%) 21.52 ( 0.05%)
Data megabytes written per second: 29.28 ( 0.00%) 29.11 (-0.58%)
Files created alone per second: 250.00 ( 0.00%) 416.00 (39.90%)
Files create/transact per second: 79.00 ( 0.00%) 76.00 (-3.95%)
Files deleted alone per second: 520.00 ( 0.00%) 420.00 (-23.81%)
Files delete/transact per second: 79.00 ( 0.00%) 76.00 (-3.95%)
MMTests Statistics: duration
User/Sys Time Running Test (seconds) 16.58 17.4
Total Elapsed Time (seconds) 218.48 222.47
VMstat Reclaim Statistics: vmscan
Direct reclaims 0 4
Direct reclaim pages scanned 0 203
Direct reclaim pages reclaimed 0 184
Kswapd pages scanned 326631 322018
Kswapd pages reclaimed 312632 309784
Kswapd low wmark quickly 1 4
Kswapd high wmark quickly 122 475
Kswapd skip congestion_wait 1 0
Pages activated 700040 705317
Pages deactivated 212113 203922
Pages written 9875 6363
Total pages scanned 326631 322221
Total pages reclaimed 312632 309968
%age total pages scanned/reclaimed 95.71% 96.20%
%age total pages scanned/written 3.02% 1.97%
proc vmstat: Faults
Major Faults 300 254
Minor Faults 645183 660284
Page ins 493588 486704
Page outs 4960088 4986704
Swap ins 1230 661
Swap outs 9869 6355
Performance is mildly affected because kswapd is no longer doing as much
work and the background memory consumer process is getting in the way.
Note that kswapd scanned and reclaimed fewer pages as it's less aggressive
and overall fewer pages were scanned and reclaimed. Swap in/out is
particularly reduced again reflecting kswapd throwing out fewer pages.
The slight performance impact is unfortunate here but it looks like a
direct result of kswapd being less aggressive. As the bug report is about
too many pages being freed by kswapd, it may have to be accepted for now.
The second test is a streaming IO benchmark that was previously used by
Johannes to show regressions in page reclaim.
MICRO
traceonly kanyzone
User/Sys Time Running Test (seconds) 29.29 28.87
Total Elapsed Time (seconds) 492.18 488.79
VMstat Reclaim Statistics: vmscan
Direct reclaims 2128 1460
Direct reclaim pages scanned 2284822 1496067
Direct reclaim pages reclaimed 148919 110937
Kswapd pages scanned 15450014 16202876
Kswapd pages reclaimed 8503697 8537897
Kswapd low wmark quickly 3100 3397
Kswapd high wmark quickly 1860 7243
Kswapd skip congestion_wait 708 801
Pages activated 9635 9573
Pages deactivated 1432 1271
Pages written 223 1130
Total pages scanned 17734836 17698943
Total pages reclaimed 8652616 8648834
%age total pages scanned/reclaimed 48.79% 48.87%
%age total pages scanned/written 0.00% 0.01%
proc vmstat: Faults
Major Faults 165 221
Minor Faults 9655785 9656506
Page ins 3880 7228
Page outs 37692940 37480076
Swap ins 0 69
Swap outs 19 15
Again fewer pages are scanned and reclaimed as expected and this time the
test completed faster. Note that kswapd is hitting its watermarks faster
(low and high wmark quickly) which I expect is due to kswapd reclaiming
fewer pages.
I also ran fs-mark, iozone and sysbench but there is nothing interesting
to report in the figures. Performance is not significantly changed and
the reclaim statistics look reasonable.
Tgis patch:
When the allocator enters its slow path, kswapd is woken up to balance the
node. It continues working until all zones within the node are balanced.
For order-0 allocations, this makes perfect sense but for higher orders it
can have unintended side-effects. If the zone sizes are imbalanced,
kswapd may reclaim heavily within a smaller zone discarding an excessive
number of pages. The user-visible behaviour is that kswapd is awake and
reclaiming even though plenty of pages are free from a suitable zone.
This patch alters the "balance" logic for high-order reclaim allowing
kswapd to stop if any suitable zone becomes balanced to reduce the number
of pages it reclaims from other zones. kswapd still tries to ensure that
order-0 watermarks for all zones are met before sleeping.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Reviewed-by: Minchan Kim <minchan.kim@gmail.com>
Reviewed-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Reviewed-by: Eric B Munson <emunson@mgebm.net>
Cc: Simon Kirby <sim@hostway.ca>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Shaohua Li <shaohua.li@intel.com>
Cc: Dave Hansen <dave@linux.vnet.ibm.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Rik van Riel <riel@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-01-14 02:46:20 +03:00
|
|
|
enum zone_type classzone_idx;
|
2012-03-23 23:56:34 +04:00
|
|
|
#ifdef CONFIG_NUMA_BALANCING
|
|
|
|
/*
|
|
|
|
* Lock serializing the per destination node AutoNUMA memory
|
|
|
|
* migration rate limiting data.
|
|
|
|
*/
|
|
|
|
spinlock_t numabalancing_migrate_lock;
|
|
|
|
|
|
|
|
/* Rate limiting time interval */
|
|
|
|
unsigned long numabalancing_migrate_next_window;
|
|
|
|
|
|
|
|
/* Number of pages migrated during the rate limiting time interval */
|
|
|
|
unsigned long numabalancing_migrate_nr_pages;
|
|
|
|
#endif
|
2005-04-17 02:20:36 +04:00
|
|
|
} pg_data_t;
|
|
|
|
|
|
|
|
#define node_present_pages(nid) (NODE_DATA(nid)->node_present_pages)
|
|
|
|
#define node_spanned_pages(nid) (NODE_DATA(nid)->node_spanned_pages)
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#ifdef CONFIG_FLAT_NODE_MEM_MAP
|
[PATCH] remove non-DISCONTIG use of pgdat->node_mem_map
This patch effectively eliminates direct use of pgdat->node_mem_map outside
of the DISCONTIG code. On a flat memory system, these fields aren't
currently used, neither are they on a sparsemem system.
There was also a node_mem_map(nid) macro on many architectures. Its use
along with the use of ->node_mem_map itself was not consistent. It has
been removed in favor of two new, more explicit, arch-independent macros:
pgdat_page_nr(pgdat, pagenr)
nid_page_nr(nid, pagenr)
I called them "pgdat" and "nid" because we overload the term "node" to mean
"NUMA node", "DISCONTIG node" or "pg_data_t" in very confusing ways. I
believe the newer names are much clearer.
These macros can be overridden in the sparsemem case with a theoretically
slower operation using node_start_pfn and pfn_to_page(), instead. We could
make this the only behavior if people want, but I don't want to change too
much at once. One thing at a time.
This patch removes more code than it adds.
Compile tested on alpha, alpha discontig, arm, arm-discontig, i386, i386
generic, NUMAQ, Summit, ppc64, ppc64 discontig, and x86_64. Full list
here: http://sr71.net/patches/2.6.12/2.6.12-rc1-mhp2/configs/
Boot tested on NUMAQ, x86 SMP and ppc64 power4/5 LPARs.
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin J. Bligh <mbligh@aracnet.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:37 +04:00
|
|
|
#define pgdat_page_nr(pgdat, pagenr) ((pgdat)->node_mem_map + (pagenr))
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#else
|
|
|
|
#define pgdat_page_nr(pgdat, pagenr) pfn_to_page((pgdat)->node_start_pfn + (pagenr))
|
|
|
|
#endif
|
[PATCH] remove non-DISCONTIG use of pgdat->node_mem_map
This patch effectively eliminates direct use of pgdat->node_mem_map outside
of the DISCONTIG code. On a flat memory system, these fields aren't
currently used, neither are they on a sparsemem system.
There was also a node_mem_map(nid) macro on many architectures. Its use
along with the use of ->node_mem_map itself was not consistent. It has
been removed in favor of two new, more explicit, arch-independent macros:
pgdat_page_nr(pgdat, pagenr)
nid_page_nr(nid, pagenr)
I called them "pgdat" and "nid" because we overload the term "node" to mean
"NUMA node", "DISCONTIG node" or "pg_data_t" in very confusing ways. I
believe the newer names are much clearer.
These macros can be overridden in the sparsemem case with a theoretically
slower operation using node_start_pfn and pfn_to_page(), instead. We could
make this the only behavior if people want, but I don't want to change too
much at once. One thing at a time.
This patch removes more code than it adds.
Compile tested on alpha, alpha discontig, arm, arm-discontig, i386, i386
generic, NUMAQ, Summit, ppc64, ppc64 discontig, and x86_64. Full list
here: http://sr71.net/patches/2.6.12/2.6.12-rc1-mhp2/configs/
Boot tested on NUMAQ, x86 SMP and ppc64 power4/5 LPARs.
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin J. Bligh <mbligh@aracnet.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:37 +04:00
|
|
|
#define nid_page_nr(nid, pagenr) pgdat_page_nr(NODE_DATA(nid),(pagenr))
|
2005-04-17 02:20:36 +04:00
|
|
|
|
2011-06-16 12:28:07 +04:00
|
|
|
#define node_start_pfn(nid) (NODE_DATA(nid)->node_start_pfn)
|
2013-02-23 04:35:27 +04:00
|
|
|
#define node_end_pfn(nid) pgdat_end_pfn(NODE_DATA(nid))
|
2011-06-16 12:28:07 +04:00
|
|
|
|
2013-02-23 04:35:27 +04:00
|
|
|
static inline unsigned long pgdat_end_pfn(pg_data_t *pgdat)
|
|
|
|
{
|
|
|
|
return pgdat->node_start_pfn + pgdat->node_spanned_pages;
|
|
|
|
}
|
|
|
|
|
|
|
|
static inline bool pgdat_is_empty(pg_data_t *pgdat)
|
|
|
|
{
|
|
|
|
return !pgdat->node_start_pfn && !pgdat->node_spanned_pages;
|
|
|
|
}
|
2011-06-16 12:28:07 +04:00
|
|
|
|
2005-10-30 04:16:52 +03:00
|
|
|
#include <linux/memory_hotplug.h>
|
|
|
|
|
2010-05-25 01:32:52 +04:00
|
|
|
extern struct mutex zonelists_mutex;
|
2012-08-01 03:43:28 +04:00
|
|
|
void build_all_zonelists(pg_data_t *pgdat, struct zone *zone);
|
mm: kswapd: stop high-order balancing when any suitable zone is balanced
Simon Kirby reported the following problem
We're seeing cases on a number of servers where cache never fully
grows to use all available memory. Sometimes we see servers with 4 GB
of memory that never seem to have less than 1.5 GB free, even with a
constantly-active VM. In some cases, these servers also swap out while
this happens, even though they are constantly reading the working set
into memory. We have been seeing this happening for a long time; I
don't think it's anything recent, and it still happens on 2.6.36.
After some debugging work by Simon, Dave Hansen and others, the prevaling
theory became that kswapd is reclaiming order-3 pages requested by SLUB
too aggressive about it.
There are two apparent problems here. On the target machine, there is a
small Normal zone in comparison to DMA32. As kswapd tries to balance all
zones, it would continually try reclaiming for Normal even though DMA32
was balanced enough for callers. The second problem is that
sleeping_prematurely() does not use the same logic as balance_pgdat() when
deciding whether to sleep or not. This keeps kswapd artifically awake.
A number of tests were run and the figures from previous postings will
look very different for a few reasons. One, the old figures were forcing
my network card to use GFP_ATOMIC in attempt to replicate Simon's problem.
Second, I previous specified slub_min_order=3 again in an attempt to
reproduce Simon's problem. In this posting, I'm depending on Simon to say
whether his problem is fixed or not and these figures are to show the
impact to the ordinary cases. Finally, the "vmscan" figures are taken
from /proc/vmstat instead of the tracepoints. There is less information
but recording is less disruptive.
The first test of relevance was postmark with a process running in the
background reading a large amount of anonymous memory in blocks. The
objective was to vaguely simulate what was happening on Simon's machine
and it's memory intensive enough to have kswapd awake.
POSTMARK
traceonly kanyzone
Transactions per second: 156.00 ( 0.00%) 153.00 (-1.96%)
Data megabytes read per second: 21.51 ( 0.00%) 21.52 ( 0.05%)
Data megabytes written per second: 29.28 ( 0.00%) 29.11 (-0.58%)
Files created alone per second: 250.00 ( 0.00%) 416.00 (39.90%)
Files create/transact per second: 79.00 ( 0.00%) 76.00 (-3.95%)
Files deleted alone per second: 520.00 ( 0.00%) 420.00 (-23.81%)
Files delete/transact per second: 79.00 ( 0.00%) 76.00 (-3.95%)
MMTests Statistics: duration
User/Sys Time Running Test (seconds) 16.58 17.4
Total Elapsed Time (seconds) 218.48 222.47
VMstat Reclaim Statistics: vmscan
Direct reclaims 0 4
Direct reclaim pages scanned 0 203
Direct reclaim pages reclaimed 0 184
Kswapd pages scanned 326631 322018
Kswapd pages reclaimed 312632 309784
Kswapd low wmark quickly 1 4
Kswapd high wmark quickly 122 475
Kswapd skip congestion_wait 1 0
Pages activated 700040 705317
Pages deactivated 212113 203922
Pages written 9875 6363
Total pages scanned 326631 322221
Total pages reclaimed 312632 309968
%age total pages scanned/reclaimed 95.71% 96.20%
%age total pages scanned/written 3.02% 1.97%
proc vmstat: Faults
Major Faults 300 254
Minor Faults 645183 660284
Page ins 493588 486704
Page outs 4960088 4986704
Swap ins 1230 661
Swap outs 9869 6355
Performance is mildly affected because kswapd is no longer doing as much
work and the background memory consumer process is getting in the way.
Note that kswapd scanned and reclaimed fewer pages as it's less aggressive
and overall fewer pages were scanned and reclaimed. Swap in/out is
particularly reduced again reflecting kswapd throwing out fewer pages.
The slight performance impact is unfortunate here but it looks like a
direct result of kswapd being less aggressive. As the bug report is about
too many pages being freed by kswapd, it may have to be accepted for now.
The second test is a streaming IO benchmark that was previously used by
Johannes to show regressions in page reclaim.
MICRO
traceonly kanyzone
User/Sys Time Running Test (seconds) 29.29 28.87
Total Elapsed Time (seconds) 492.18 488.79
VMstat Reclaim Statistics: vmscan
Direct reclaims 2128 1460
Direct reclaim pages scanned 2284822 1496067
Direct reclaim pages reclaimed 148919 110937
Kswapd pages scanned 15450014 16202876
Kswapd pages reclaimed 8503697 8537897
Kswapd low wmark quickly 3100 3397
Kswapd high wmark quickly 1860 7243
Kswapd skip congestion_wait 708 801
Pages activated 9635 9573
Pages deactivated 1432 1271
Pages written 223 1130
Total pages scanned 17734836 17698943
Total pages reclaimed 8652616 8648834
%age total pages scanned/reclaimed 48.79% 48.87%
%age total pages scanned/written 0.00% 0.01%
proc vmstat: Faults
Major Faults 165 221
Minor Faults 9655785 9656506
Page ins 3880 7228
Page outs 37692940 37480076
Swap ins 0 69
Swap outs 19 15
Again fewer pages are scanned and reclaimed as expected and this time the
test completed faster. Note that kswapd is hitting its watermarks faster
(low and high wmark quickly) which I expect is due to kswapd reclaiming
fewer pages.
I also ran fs-mark, iozone and sysbench but there is nothing interesting
to report in the figures. Performance is not significantly changed and
the reclaim statistics look reasonable.
Tgis patch:
When the allocator enters its slow path, kswapd is woken up to balance the
node. It continues working until all zones within the node are balanced.
For order-0 allocations, this makes perfect sense but for higher orders it
can have unintended side-effects. If the zone sizes are imbalanced,
kswapd may reclaim heavily within a smaller zone discarding an excessive
number of pages. The user-visible behaviour is that kswapd is awake and
reclaiming even though plenty of pages are free from a suitable zone.
This patch alters the "balance" logic for high-order reclaim allowing
kswapd to stop if any suitable zone becomes balanced to reduce the number
of pages it reclaims from other zones. kswapd still tries to ensure that
order-0 watermarks for all zones are met before sleeping.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Reviewed-by: Minchan Kim <minchan.kim@gmail.com>
Reviewed-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Reviewed-by: Eric B Munson <emunson@mgebm.net>
Cc: Simon Kirby <sim@hostway.ca>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Shaohua Li <shaohua.li@intel.com>
Cc: Dave Hansen <dave@linux.vnet.ibm.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Rik van Riel <riel@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-01-14 02:46:20 +03:00
|
|
|
void wakeup_kswapd(struct zone *zone, int order, enum zone_type classzone_idx);
|
mm: page allocator: adjust the per-cpu counter threshold when memory is low
Commit aa45484 ("calculate a better estimate of NR_FREE_PAGES when memory
is low") noted that watermarks were based on the vmstat NR_FREE_PAGES. To
avoid synchronization overhead, these counters are maintained on a per-cpu
basis and drained both periodically and when a threshold is above a
threshold. On large CPU systems, the difference between the estimate and
real value of NR_FREE_PAGES can be very high. The system can get into a
case where pages are allocated far below the min watermark potentially
causing livelock issues. The commit solved the problem by taking a better
reading of NR_FREE_PAGES when memory was low.
Unfortately, as reported by Shaohua Li this accurate reading can consume a
large amount of CPU time on systems with many sockets due to cache line
bouncing. This patch takes a different approach. For large machines
where counter drift might be unsafe and while kswapd is awake, the per-cpu
thresholds for the target pgdat are reduced to limit the level of drift to
what should be a safe level. This incurs a performance penalty in heavy
memory pressure by a factor that depends on the workload and the machine
but the machine should function correctly without accidentally exhausting
all memory on a node. There is an additional cost when kswapd wakes and
sleeps but the event is not expected to be frequent - in Shaohua's test
case, there was one recorded sleep and wake event at least.
To ensure that kswapd wakes up, a safe version of zone_watermark_ok() is
introduced that takes a more accurate reading of NR_FREE_PAGES when called
from wakeup_kswapd, when deciding whether it is really safe to go back to
sleep in sleeping_prematurely() and when deciding if a zone is really
balanced or not in balance_pgdat(). We are still using an expensive
function but limiting how often it is called.
When the test case is reproduced, the time spent in the watermark
functions is reduced. The following report is on the percentage of time
spent cumulatively spent in the functions zone_nr_free_pages(),
zone_watermark_ok(), __zone_watermark_ok(), zone_watermark_ok_safe(),
zone_page_state_snapshot(), zone_page_state().
vanilla 11.6615%
disable-threshold 0.2584%
David said:
: We had to pull aa454840 "mm: page allocator: calculate a better estimate
: of NR_FREE_PAGES when memory is low and kswapd is awake" from 2.6.36
: internally because tests showed that it would cause the machine to stall
: as the result of heavy kswapd activity. I merged it back with this fix as
: it is pending in the -mm tree and it solves the issue we were seeing, so I
: definitely think this should be pushed to -stable (and I would seriously
: consider it for 2.6.37 inclusion even at this late date).
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Reported-by: Shaohua Li <shaohua.li@intel.com>
Reviewed-by: Christoph Lameter <cl@linux.com>
Tested-by: Nicolas Bareil <nico@chdir.org>
Cc: David Rientjes <rientjes@google.com>
Cc: Kyle McMartin <kyle@mcmartin.ca>
Cc: <stable@kernel.org> [2.6.37.1, 2.6.36.x]
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-01-14 02:45:41 +03:00
|
|
|
bool zone_watermark_ok(struct zone *z, int order, unsigned long mark,
|
|
|
|
int classzone_idx, int alloc_flags);
|
|
|
|
bool zone_watermark_ok_safe(struct zone *z, int order, unsigned long mark,
|
2005-11-14 03:06:43 +03:00
|
|
|
int classzone_idx, int alloc_flags);
|
2007-01-11 10:15:30 +03:00
|
|
|
enum memmap_context {
|
|
|
|
MEMMAP_EARLY,
|
|
|
|
MEMMAP_HOTPLUG,
|
|
|
|
};
|
2006-06-23 13:03:10 +04:00
|
|
|
extern int init_currently_empty_zone(struct zone *zone, unsigned long start_pfn,
|
2007-01-11 10:15:30 +03:00
|
|
|
unsigned long size,
|
|
|
|
enum memmap_context context);
|
2006-06-23 13:03:10 +04:00
|
|
|
|
memcg: fix hotplugged memory zone oops
When MEMCG is configured on (even when it's disabled by boot option),
when adding or removing a page to/from its lru list, the zone pointer
used for stats updates is nowadays taken from the struct lruvec. (On
many configurations, calculating zone from page is slower.)
But we have no code to update all the lruvecs (per zone, per memcg) when
a memory node is hotadded. Here's an extract from the oops which
results when running numactl to bind a program to a newly onlined node:
BUG: unable to handle kernel NULL pointer dereference at 0000000000000f60
IP: __mod_zone_page_state+0x9/0x60
Pid: 1219, comm: numactl Not tainted 3.6.0-rc5+ #180 Bochs Bochs
Process numactl (pid: 1219, threadinfo ffff880039abc000, task ffff8800383c4ce0)
Call Trace:
__pagevec_lru_add_fn+0xdf/0x140
pagevec_lru_move_fn+0xb1/0x100
__pagevec_lru_add+0x1c/0x30
lru_add_drain_cpu+0xa3/0x130
lru_add_drain+0x2f/0x40
...
The natural solution might be to use a memcg callback whenever memory is
hotadded; but that solution has not been scoped out, and it happens that
we do have an easy location at which to update lruvec->zone. The lruvec
pointer is discovered either by mem_cgroup_zone_lruvec() or by
mem_cgroup_page_lruvec(), and both of those do know the right zone.
So check and set lruvec->zone in those; and remove the inadequate
attempt to set lruvec->zone from lruvec_init(), which is called before
NODE_DATA(node) has been allocated in such cases.
Ah, there was one exceptionr. For no particularly good reason,
mem_cgroup_force_empty_list() has its own code for deciding lruvec.
Change it to use the standard mem_cgroup_zone_lruvec() and
mem_cgroup_get_lru_size() too. In fact it was already safe against such
an oops (the lru lists in danger could only be empty), but we're better
proofed against future changes this way.
I've marked this for stable (3.6) since we introduced the problem in 3.5
(now closed to stable); but I have no idea if this is the only fix
needed to get memory hotadd working with memcg in 3.6, and received no
answer when I enquired twice before.
Reported-by: Tang Chen <tangchen@cn.fujitsu.com>
Signed-off-by: Hugh Dickins <hughd@google.com>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Konstantin Khlebnikov <khlebnikov@openvz.org>
Cc: Wen Congyang <wency@cn.fujitsu.com>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-11-17 02:14:54 +04:00
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extern void lruvec_init(struct lruvec *lruvec);
|
2012-05-30 02:06:58 +04:00
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static inline struct zone *lruvec_zone(struct lruvec *lruvec)
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{
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2012-08-01 03:43:02 +04:00
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#ifdef CONFIG_MEMCG
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2012-05-30 02:06:58 +04:00
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return lruvec->zone;
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#else
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return container_of(lruvec, struct zone, lruvec);
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#endif
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}
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2005-04-17 02:20:36 +04:00
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#ifdef CONFIG_HAVE_MEMORY_PRESENT
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void memory_present(int nid, unsigned long start, unsigned long end);
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#else
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static inline void memory_present(int nid, unsigned long start, unsigned long end) {}
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#endif
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2010-05-27 01:45:00 +04:00
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#ifdef CONFIG_HAVE_MEMORYLESS_NODES
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int local_memory_node(int node_id);
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#else
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static inline int local_memory_node(int node_id) { return node_id; };
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#endif
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2005-04-17 02:20:36 +04:00
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#ifdef CONFIG_NEED_NODE_MEMMAP_SIZE
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unsigned long __init node_memmap_size_bytes(int, unsigned long, unsigned long);
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#endif
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/*
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* zone_idx() returns 0 for the ZONE_DMA zone, 1 for the ZONE_NORMAL zone, etc.
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*/
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#define zone_idx(zone) ((zone) - (zone)->zone_pgdat->node_zones)
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2006-01-06 11:11:15 +03:00
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static inline int populated_zone(struct zone *zone)
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{
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return (!!zone->present_pages);
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}
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2007-07-17 15:03:12 +04:00
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extern int movable_zone;
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static inline int zone_movable_is_highmem(void)
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{
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2012-08-01 03:43:14 +04:00
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#if defined(CONFIG_HIGHMEM) && defined(CONFIG_HAVE_MEMBLOCK_NODE_MAP)
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2007-07-17 15:03:12 +04:00
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return movable_zone == ZONE_HIGHMEM;
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#else
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return 0;
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#endif
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}
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2006-09-26 10:31:13 +04:00
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static inline int is_highmem_idx(enum zone_type idx)
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2005-04-17 02:20:36 +04:00
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{
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2006-09-26 10:31:14 +04:00
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#ifdef CONFIG_HIGHMEM
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2007-07-17 15:03:12 +04:00
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return (idx == ZONE_HIGHMEM ||
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(idx == ZONE_MOVABLE && zone_movable_is_highmem()));
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2006-09-26 10:31:14 +04:00
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#else
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return 0;
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#endif
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2005-04-17 02:20:36 +04:00
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}
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/**
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* is_highmem - helper function to quickly check if a struct zone is a
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* highmem zone or not. This is an attempt to keep references
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* to ZONE_{DMA/NORMAL/HIGHMEM/etc} in general code to a minimum.
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* @zone - pointer to struct zone variable
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*/
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static inline int is_highmem(struct zone *zone)
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{
|
2006-09-26 10:31:14 +04:00
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#ifdef CONFIG_HIGHMEM
|
remove sparse warning for mmzone.h
include/linux/mmzone.h:640:22: warning: potentially expensive pointer subtraction
Calculate the offset into the node_zones array rather than the index
using casts to (char *) and comparing against the index * sizeof(struct zone).
On X86_32 this saves a sar, but code size increases by one byte per
is_highmem() use due to 32-bit cmps rather than 16 bit cmps.
Before:
207: 2b 80 8c 07 00 00 sub 0x78c(%eax),%eax
20d: c1 f8 0b sar $0xb,%eax
210: 83 f8 02 cmp $0x2,%eax
213: 74 16 je 22b <kmap_atomic_prot+0x144>
215: 83 f8 03 cmp $0x3,%eax
218: 0f 85 8f 00 00 00 jne 2ad <kmap_atomic_prot+0x1c6>
21e: 83 3d 00 00 00 00 02 cmpl $0x2,0x0
225: 0f 85 82 00 00 00 jne 2ad <kmap_atomic_prot+0x1c6>
22b: 64 a1 00 00 00 00 mov %fs:0x0,%eax
After:
207: 2b 80 8c 07 00 00 sub 0x78c(%eax),%eax
20d: 3d 00 10 00 00 cmp $0x1000,%eax
212: 74 18 je 22c <kmap_atomic_prot+0x145>
214: 3d 00 18 00 00 cmp $0x1800,%eax
219: 0f 85 8f 00 00 00 jne 2ae <kmap_atomic_prot+0x1c7>
21f: 83 3d 00 00 00 00 02 cmpl $0x2,0x0
226: 0f 85 82 00 00 00 jne 2ae <kmap_atomic_prot+0x1c7>
22c: 64 a1 00 00 00 00 mov %fs:0x0,%eax
[akpm@linux-foundation.org: coding-style fixes]
Signed-off-by: Harvey Harrison <harvey.harrison@gmail.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-04-28 13:12:07 +04:00
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int zone_off = (char *)zone - (char *)zone->zone_pgdat->node_zones;
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return zone_off == ZONE_HIGHMEM * sizeof(*zone) ||
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(zone_off == ZONE_MOVABLE * sizeof(*zone) &&
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zone_movable_is_highmem());
|
2006-09-26 10:31:14 +04:00
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#else
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return 0;
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#endif
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2005-04-17 02:20:36 +04:00
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}
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/* These two functions are used to setup the per zone pages min values */
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struct ctl_table;
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2009-09-24 02:57:19 +04:00
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int min_free_kbytes_sysctl_handler(struct ctl_table *, int,
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2005-04-17 02:20:36 +04:00
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void __user *, size_t *, loff_t *);
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extern int sysctl_lowmem_reserve_ratio[MAX_NR_ZONES-1];
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2009-09-24 02:57:19 +04:00
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int lowmem_reserve_ratio_sysctl_handler(struct ctl_table *, int,
|
2005-04-17 02:20:36 +04:00
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void __user *, size_t *, loff_t *);
|
2009-09-24 02:57:19 +04:00
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int percpu_pagelist_fraction_sysctl_handler(struct ctl_table *, int,
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2006-01-08 12:00:40 +03:00
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void __user *, size_t *, loff_t *);
|
2006-07-03 11:24:13 +04:00
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int sysctl_min_unmapped_ratio_sysctl_handler(struct ctl_table *, int,
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2009-09-24 02:57:19 +04:00
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void __user *, size_t *, loff_t *);
|
2006-09-26 10:31:52 +04:00
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int sysctl_min_slab_ratio_sysctl_handler(struct ctl_table *, int,
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2009-09-24 02:57:19 +04:00
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void __user *, size_t *, loff_t *);
|
2005-04-17 02:20:36 +04:00
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2007-07-16 10:38:01 +04:00
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extern int numa_zonelist_order_handler(struct ctl_table *, int,
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2009-09-24 02:57:19 +04:00
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void __user *, size_t *, loff_t *);
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2007-07-16 10:38:01 +04:00
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extern char numa_zonelist_order[];
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#define NUMA_ZONELIST_ORDER_LEN 16 /* string buffer size */
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2005-06-23 11:07:47 +04:00
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#ifndef CONFIG_NEED_MULTIPLE_NODES
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2005-04-17 02:20:36 +04:00
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extern struct pglist_data contig_page_data;
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#define NODE_DATA(nid) (&contig_page_data)
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#define NODE_MEM_MAP(nid) mem_map
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|
2005-06-23 11:07:47 +04:00
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#else /* CONFIG_NEED_MULTIPLE_NODES */
|
2005-04-17 02:20:36 +04:00
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#include <asm/mmzone.h>
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2005-06-23 11:07:47 +04:00
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#endif /* !CONFIG_NEED_MULTIPLE_NODES */
|
2005-06-23 11:07:40 +04:00
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|
2006-03-27 13:16:02 +04:00
|
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|
extern struct pglist_data *first_online_pgdat(void);
|
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extern struct pglist_data *next_online_pgdat(struct pglist_data *pgdat);
|
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extern struct zone *next_zone(struct zone *zone);
|
2006-03-27 13:15:57 +04:00
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|
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/**
|
2008-05-24 00:05:01 +04:00
|
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|
* for_each_online_pgdat - helper macro to iterate over all online nodes
|
2006-03-27 13:15:57 +04:00
|
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|
* @pgdat - pointer to a pg_data_t variable
|
|
|
|
*/
|
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|
#define for_each_online_pgdat(pgdat) \
|
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|
for (pgdat = first_online_pgdat(); \
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|
pgdat; \
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|
pgdat = next_online_pgdat(pgdat))
|
|
|
|
/**
|
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|
|
* for_each_zone - helper macro to iterate over all memory zones
|
|
|
|
* @zone - pointer to struct zone variable
|
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|
*
|
|
|
|
* The user only needs to declare the zone variable, for_each_zone
|
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|
* fills it in.
|
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|
*/
|
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|
#define for_each_zone(zone) \
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for (zone = (first_online_pgdat())->node_zones; \
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zone; \
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zone = next_zone(zone))
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|
2009-04-01 02:19:31 +04:00
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|
#define for_each_populated_zone(zone) \
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for (zone = (first_online_pgdat())->node_zones; \
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zone; \
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zone = next_zone(zone)) \
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|
if (!populated_zone(zone)) \
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; /* do nothing */ \
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else
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|
2008-04-28 13:12:17 +04:00
|
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|
static inline struct zone *zonelist_zone(struct zoneref *zoneref)
|
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|
{
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|
return zoneref->zone;
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}
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static inline int zonelist_zone_idx(struct zoneref *zoneref)
|
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|
|
{
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|
return zoneref->zone_idx;
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}
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static inline int zonelist_node_idx(struct zoneref *zoneref)
|
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|
{
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|
#ifdef CONFIG_NUMA
|
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|
|
/* zone_to_nid not available in this context */
|
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|
return zoneref->zone->node;
|
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#else
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return 0;
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#endif /* CONFIG_NUMA */
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}
|
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|
2008-04-28 13:12:18 +04:00
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/**
|
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|
* next_zones_zonelist - Returns the next zone at or below highest_zoneidx within the allowed nodemask using a cursor within a zonelist as a starting point
|
|
|
|
* @z - The cursor used as a starting point for the search
|
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|
* @highest_zoneidx - The zone index of the highest zone to return
|
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|
* @nodes - An optional nodemask to filter the zonelist with
|
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|
* @zone - The first suitable zone found is returned via this parameter
|
|
|
|
*
|
|
|
|
* This function returns the next zone at or below a given zone index that is
|
|
|
|
* within the allowed nodemask using a cursor as the starting point for the
|
2008-09-13 13:33:19 +04:00
|
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|
* search. The zoneref returned is a cursor that represents the current zone
|
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|
* being examined. It should be advanced by one before calling
|
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|
|
* next_zones_zonelist again.
|
2008-04-28 13:12:18 +04:00
|
|
|
*/
|
|
|
|
struct zoneref *next_zones_zonelist(struct zoneref *z,
|
|
|
|
enum zone_type highest_zoneidx,
|
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|
|
nodemask_t *nodes,
|
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|
|
struct zone **zone);
|
2008-04-28 13:12:17 +04:00
|
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|
2008-04-28 13:12:18 +04:00
|
|
|
/**
|
|
|
|
* first_zones_zonelist - Returns the first zone at or below highest_zoneidx within the allowed nodemask in a zonelist
|
|
|
|
* @zonelist - The zonelist to search for a suitable zone
|
|
|
|
* @highest_zoneidx - The zone index of the highest zone to return
|
|
|
|
* @nodes - An optional nodemask to filter the zonelist with
|
|
|
|
* @zone - The first suitable zone found is returned via this parameter
|
|
|
|
*
|
|
|
|
* This function returns the first zone at or below a given zone index that is
|
|
|
|
* within the allowed nodemask. The zoneref returned is a cursor that can be
|
2008-09-13 13:33:19 +04:00
|
|
|
* used to iterate the zonelist with next_zones_zonelist by advancing it by
|
|
|
|
* one before calling.
|
2008-04-28 13:12:18 +04:00
|
|
|
*/
|
2008-04-28 13:12:17 +04:00
|
|
|
static inline struct zoneref *first_zones_zonelist(struct zonelist *zonelist,
|
2008-04-28 13:12:18 +04:00
|
|
|
enum zone_type highest_zoneidx,
|
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|
|
nodemask_t *nodes,
|
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|
|
struct zone **zone)
|
2008-04-28 13:12:16 +04:00
|
|
|
{
|
2008-04-28 13:12:18 +04:00
|
|
|
return next_zones_zonelist(zonelist->_zonerefs, highest_zoneidx, nodes,
|
|
|
|
zone);
|
2008-04-28 13:12:16 +04:00
|
|
|
}
|
|
|
|
|
2008-04-28 13:12:18 +04:00
|
|
|
/**
|
|
|
|
* for_each_zone_zonelist_nodemask - helper macro to iterate over valid zones in a zonelist at or below a given zone index and within a nodemask
|
|
|
|
* @zone - The current zone in the iterator
|
|
|
|
* @z - The current pointer within zonelist->zones being iterated
|
|
|
|
* @zlist - The zonelist being iterated
|
|
|
|
* @highidx - The zone index of the highest zone to return
|
|
|
|
* @nodemask - Nodemask allowed by the allocator
|
|
|
|
*
|
|
|
|
* This iterator iterates though all zones at or below a given zone index and
|
|
|
|
* within a given nodemask
|
|
|
|
*/
|
|
|
|
#define for_each_zone_zonelist_nodemask(zone, z, zlist, highidx, nodemask) \
|
|
|
|
for (z = first_zones_zonelist(zlist, highidx, nodemask, &zone); \
|
|
|
|
zone; \
|
2008-09-13 13:33:19 +04:00
|
|
|
z = next_zones_zonelist(++z, highidx, nodemask, &zone)) \
|
2008-04-28 13:12:16 +04:00
|
|
|
|
|
|
|
/**
|
|
|
|
* for_each_zone_zonelist - helper macro to iterate over valid zones in a zonelist at or below a given zone index
|
|
|
|
* @zone - The current zone in the iterator
|
|
|
|
* @z - The current pointer within zonelist->zones being iterated
|
|
|
|
* @zlist - The zonelist being iterated
|
|
|
|
* @highidx - The zone index of the highest zone to return
|
|
|
|
*
|
|
|
|
* This iterator iterates though all zones at or below a given zone index.
|
|
|
|
*/
|
|
|
|
#define for_each_zone_zonelist(zone, z, zlist, highidx) \
|
2008-04-28 13:12:18 +04:00
|
|
|
for_each_zone_zonelist_nodemask(zone, z, zlist, highidx, NULL)
|
2008-04-28 13:12:16 +04:00
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#ifdef CONFIG_SPARSEMEM
|
|
|
|
#include <asm/sparsemem.h>
|
|
|
|
#endif
|
|
|
|
|
[PATCH] Introduce mechanism for registering active regions of memory
At a basic level, architectures define structures to record where active
ranges of page frames are located. Once located, the code to calculate zone
sizes and holes in each architecture is very similar. Some of this zone and
hole sizing code is difficult to read for no good reason. This set of patches
eliminates the similar-looking architecture-specific code.
The patches introduce a mechanism where architectures register where the
active ranges of page frames are with add_active_range(). When all areas have
been discovered, free_area_init_nodes() is called to initialise the pgdat and
zones. The zone sizes and holes are then calculated in an architecture
independent manner.
Patch 1 introduces the mechanism for registering and initialising PFN ranges
Patch 2 changes ppc to use the mechanism - 139 arch-specific LOC removed
Patch 3 changes x86 to use the mechanism - 136 arch-specific LOC removed
Patch 4 changes x86_64 to use the mechanism - 74 arch-specific LOC removed
Patch 5 changes ia64 to use the mechanism - 52 arch-specific LOC removed
Patch 6 accounts for mem_map as a memory hole as the pages are not reclaimable.
It adjusts the watermarks slightly
Tony Luck has successfully tested for ia64 on Itanium with tiger_defconfig,
gensparse_defconfig and defconfig. Bob Picco has also tested and debugged on
IA64. Jack Steiner successfully boot tested on a mammoth SGI IA64-based
machine. These were on patches against 2.6.17-rc1 and release 3 of these
patches but there have been no ia64-changes since release 3.
There are differences in the zone sizes for x86_64 as the arch-specific code
for x86_64 accounts the kernel image and the starting mem_maps as memory holes
but the architecture-independent code accounts the memory as present.
The big benefit of this set of patches is a sizable reduction of
architecture-specific code, some of which is very hairy. There should be a
greater reduction when other architectures use the same mechanisms for zone
and hole sizing but I lack the hardware to test on.
Additional credit;
Dave Hansen for the initial suggestion and comments on early patches
Andy Whitcroft for reviewing early versions and catching numerous
errors
Tony Luck for testing and debugging on IA64
Bob Picco for fixing bugs related to pfn registration, reviewing a
number of patch revisions, providing a number of suggestions
on future direction and testing heavily
Jack Steiner and Robin Holt for testing on IA64 and clarifying
issues related to memory holes
Yasunori for testing on IA64
Andi Kleen for reviewing and feeding back about x86_64
Christian Kujau for providing valuable information related to ACPI
problems on x86_64 and testing potential fixes
This patch:
Define the structure to represent an active range of page frames within a node
in an architecture independent manner. Architectures are expected to register
active ranges of PFNs using add_active_range(nid, start_pfn, end_pfn) and call
free_area_init_nodes() passing the PFNs of the end of each zone.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Andy Whitcroft <apw@shadowen.org>
Cc: Andi Kleen <ak@muc.de>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Paul Mackerras <paulus@samba.org>
Cc: "Keith Mannthey" <kmannth@gmail.com>
Cc: "Luck, Tony" <tony.luck@intel.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-09-27 12:49:43 +04:00
|
|
|
#if !defined(CONFIG_HAVE_ARCH_EARLY_PFN_TO_NID) && \
|
2011-12-08 22:22:09 +04:00
|
|
|
!defined(CONFIG_HAVE_MEMBLOCK_NODE_MAP)
|
2008-04-28 13:12:39 +04:00
|
|
|
static inline unsigned long early_pfn_to_nid(unsigned long pfn)
|
|
|
|
{
|
|
|
|
return 0;
|
|
|
|
}
|
2005-06-23 11:07:52 +04:00
|
|
|
#endif
|
|
|
|
|
2006-01-06 11:10:53 +03:00
|
|
|
#ifdef CONFIG_FLATMEM
|
|
|
|
#define pfn_to_nid(pfn) (0)
|
|
|
|
#endif
|
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#ifdef CONFIG_SPARSEMEM
|
|
|
|
|
|
|
|
/*
|
|
|
|
* SECTION_SHIFT #bits space required to store a section #
|
|
|
|
*
|
|
|
|
* PA_SECTION_SHIFT physical address to/from section number
|
|
|
|
* PFN_SECTION_SHIFT pfn to/from section number
|
|
|
|
*/
|
|
|
|
#define PA_SECTION_SHIFT (SECTION_SIZE_BITS)
|
|
|
|
#define PFN_SECTION_SHIFT (SECTION_SIZE_BITS - PAGE_SHIFT)
|
|
|
|
|
|
|
|
#define NR_MEM_SECTIONS (1UL << SECTIONS_SHIFT)
|
|
|
|
|
|
|
|
#define PAGES_PER_SECTION (1UL << PFN_SECTION_SHIFT)
|
|
|
|
#define PAGE_SECTION_MASK (~(PAGES_PER_SECTION-1))
|
|
|
|
|
Add a bitmap that is used to track flags affecting a block of pages
Here is the latest revision of the anti-fragmentation patches. Of particular
note in this version is special treatment of high-order atomic allocations.
Care is taken to group them together and avoid grouping pages of other types
near them. Artifical tests imply that it works. I'm trying to get the
hardware together that would allow setting up of a "real" test. If anyone
already has a setup and test that can trigger the atomic-allocation problem,
I'd appreciate a test of these patches and a report. The second major change
is that these patches will apply cleanly with patches that implement
anti-fragmentation through zones.
kernbench shows effectively no performance difference varying between -0.2%
and +2% on a variety of test machines. Success rates for huge page allocation
are dramatically increased. For example, on a ppc64 machine, the vanilla
kernel was only able to allocate 1% of memory as a hugepage and this was due
to a single hugepage reserved as min_free_kbytes. With these patches applied,
17% was allocatable as superpages. With reclaim-related fixes from Andy
Whitcroft, it was 40% and further reclaim-related improvements should increase
this further.
Changelog Since V28
o Group high-order atomic allocations together
o It is no longer required to set min_free_kbytes to 10% of memory. A value
of 16384 in most cases will be sufficient
o Now applied with zone-based anti-fragmentation
o Fix incorrect VM_BUG_ON within buffered_rmqueue()
o Reorder the stack so later patches do not back out work from earlier patches
o Fix bug were journal pages were being treated as movable
o Bias placement of non-movable pages to lower PFNs
o More agressive clustering of reclaimable pages in reactions to workloads
like updatedb that flood the size of inode caches
Changelog Since V27
o Renamed anti-fragmentation to Page Clustering. Anti-fragmentation was giving
the mistaken impression that it was the 100% solution for high order
allocations. Instead, it greatly increases the chances high-order
allocations will succeed and lays the foundation for defragmentation and
memory hot-remove to work properly
o Redefine page groupings based on ability to migrate or reclaim instead of
basing on reclaimability alone
o Get rid of spurious inits
o Per-cpu lists are no longer split up per-type. Instead the per-cpu list is
searched for a page of the appropriate type
o Added more explanation commentary
o Fix up bug in pageblock code where bitmap was used before being initalised
Changelog Since V26
o Fix double init of lists in setup_pageset
Changelog Since V25
o Fix loop order of for_each_rclmtype_order so that order of loop matches args
o gfpflags_to_rclmtype uses gfp_t instead of unsigned long
o Rename get_pageblock_type() to get_page_rclmtype()
o Fix alignment problem in move_freepages()
o Add mechanism for assigning flags to blocks of pages instead of page->flags
o On fallback, do not examine the preferred list of free pages a second time
The purpose of these patches is to reduce external fragmentation by grouping
pages of related types together. When pages are migrated (or reclaimed under
memory pressure), large contiguous pages will be freed.
This patch works by categorising allocations by their ability to migrate;
Movable - The pages may be moved with the page migration mechanism. These are
generally userspace pages.
Reclaimable - These are allocations for some kernel caches that are
reclaimable or allocations that are known to be very short-lived.
Unmovable - These are pages that are allocated by the kernel that
are not trivially reclaimed. For example, the memory allocated for a
loaded module would be in this category. By default, allocations are
considered to be of this type
HighAtomic - These are high-order allocations belonging to callers that
cannot sleep or perform any IO. In practice, this is restricted to
jumbo frame allocation for network receive. It is assumed that the
allocations are short-lived
Instead of having one MAX_ORDER-sized array of free lists in struct free_area,
there is one for each type of reclaimability. Once a 2^MAX_ORDER block of
pages is split for a type of allocation, it is added to the free-lists for
that type, in effect reserving it. Hence, over time, pages of the different
types can be clustered together.
When the preferred freelists are expired, the largest possible block is taken
from an alternative list. Buddies that are split from that large block are
placed on the preferred allocation-type freelists to mitigate fragmentation.
This implementation gives best-effort for low fragmentation in all zones.
Ideally, min_free_kbytes needs to be set to a value equal to 4 * (1 <<
(MAX_ORDER-1)) pages in most cases. This would be 16384 on x86 and x86_64 for
example.
Our tests show that about 60-70% of physical memory can be allocated on a
desktop after a few days uptime. In benchmarks and stress tests, we are
finding that 80% of memory is available as contiguous blocks at the end of the
test. To compare, a standard kernel was getting < 1% of memory as large pages
on a desktop and about 8-12% of memory as large pages at the end of stress
tests.
Following this email are 12 patches that implement thie page grouping feature.
The first patch introduces a mechanism for storing flags related to a whole
block of pages. Then allocations are split between movable and all other
allocations. Following that are patches to deal with per-cpu pages and make
the mechanism configurable. The next patch moves free pages between lists
when partially allocated blocks are used for pages of another migrate type.
The second last patch groups reclaimable kernel allocations such as inode
caches together. The final patch related to groupings keeps high-order atomic
allocations.
The last two patches are more concerned with control of fragmentation. The
second last patch biases placement of non-movable allocations towards the
start of memory. This is with a view of supporting memory hot-remove of DIMMs
with higher PFNs in the future. The biasing could be enforced a lot heavier
but it would cost. The last patch agressively clusters reclaimable pages like
inode caches together.
The fragmentation reduction strategy needs to track if pages within a block
can be moved or reclaimed so that pages are freed to the appropriate list.
This patch adds a bitmap for flags affecting a whole a MAX_ORDER block of
pages.
In non-SPARSEMEM configurations, the bitmap is stored in the struct zone and
allocated during initialisation. SPARSEMEM statically allocates the bitmap in
a struct mem_section so that bitmaps do not have to be resized during memory
hotadd. This wastes a small amount of memory per unused section (usually
sizeof(unsigned long)) but the complexity of dynamically allocating the memory
is quite high.
Additional credit to Andy Whitcroft who reviewed up an earlier implementation
of the mechanism an suggested how to make it a *lot* cleaner.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Cc: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:25:47 +04:00
|
|
|
#define SECTION_BLOCKFLAGS_BITS \
|
2007-10-16 12:26:01 +04:00
|
|
|
((1UL << (PFN_SECTION_SHIFT - pageblock_order)) * NR_PAGEBLOCK_BITS)
|
Add a bitmap that is used to track flags affecting a block of pages
Here is the latest revision of the anti-fragmentation patches. Of particular
note in this version is special treatment of high-order atomic allocations.
Care is taken to group them together and avoid grouping pages of other types
near them. Artifical tests imply that it works. I'm trying to get the
hardware together that would allow setting up of a "real" test. If anyone
already has a setup and test that can trigger the atomic-allocation problem,
I'd appreciate a test of these patches and a report. The second major change
is that these patches will apply cleanly with patches that implement
anti-fragmentation through zones.
kernbench shows effectively no performance difference varying between -0.2%
and +2% on a variety of test machines. Success rates for huge page allocation
are dramatically increased. For example, on a ppc64 machine, the vanilla
kernel was only able to allocate 1% of memory as a hugepage and this was due
to a single hugepage reserved as min_free_kbytes. With these patches applied,
17% was allocatable as superpages. With reclaim-related fixes from Andy
Whitcroft, it was 40% and further reclaim-related improvements should increase
this further.
Changelog Since V28
o Group high-order atomic allocations together
o It is no longer required to set min_free_kbytes to 10% of memory. A value
of 16384 in most cases will be sufficient
o Now applied with zone-based anti-fragmentation
o Fix incorrect VM_BUG_ON within buffered_rmqueue()
o Reorder the stack so later patches do not back out work from earlier patches
o Fix bug were journal pages were being treated as movable
o Bias placement of non-movable pages to lower PFNs
o More agressive clustering of reclaimable pages in reactions to workloads
like updatedb that flood the size of inode caches
Changelog Since V27
o Renamed anti-fragmentation to Page Clustering. Anti-fragmentation was giving
the mistaken impression that it was the 100% solution for high order
allocations. Instead, it greatly increases the chances high-order
allocations will succeed and lays the foundation for defragmentation and
memory hot-remove to work properly
o Redefine page groupings based on ability to migrate or reclaim instead of
basing on reclaimability alone
o Get rid of spurious inits
o Per-cpu lists are no longer split up per-type. Instead the per-cpu list is
searched for a page of the appropriate type
o Added more explanation commentary
o Fix up bug in pageblock code where bitmap was used before being initalised
Changelog Since V26
o Fix double init of lists in setup_pageset
Changelog Since V25
o Fix loop order of for_each_rclmtype_order so that order of loop matches args
o gfpflags_to_rclmtype uses gfp_t instead of unsigned long
o Rename get_pageblock_type() to get_page_rclmtype()
o Fix alignment problem in move_freepages()
o Add mechanism for assigning flags to blocks of pages instead of page->flags
o On fallback, do not examine the preferred list of free pages a second time
The purpose of these patches is to reduce external fragmentation by grouping
pages of related types together. When pages are migrated (or reclaimed under
memory pressure), large contiguous pages will be freed.
This patch works by categorising allocations by their ability to migrate;
Movable - The pages may be moved with the page migration mechanism. These are
generally userspace pages.
Reclaimable - These are allocations for some kernel caches that are
reclaimable or allocations that are known to be very short-lived.
Unmovable - These are pages that are allocated by the kernel that
are not trivially reclaimed. For example, the memory allocated for a
loaded module would be in this category. By default, allocations are
considered to be of this type
HighAtomic - These are high-order allocations belonging to callers that
cannot sleep or perform any IO. In practice, this is restricted to
jumbo frame allocation for network receive. It is assumed that the
allocations are short-lived
Instead of having one MAX_ORDER-sized array of free lists in struct free_area,
there is one for each type of reclaimability. Once a 2^MAX_ORDER block of
pages is split for a type of allocation, it is added to the free-lists for
that type, in effect reserving it. Hence, over time, pages of the different
types can be clustered together.
When the preferred freelists are expired, the largest possible block is taken
from an alternative list. Buddies that are split from that large block are
placed on the preferred allocation-type freelists to mitigate fragmentation.
This implementation gives best-effort for low fragmentation in all zones.
Ideally, min_free_kbytes needs to be set to a value equal to 4 * (1 <<
(MAX_ORDER-1)) pages in most cases. This would be 16384 on x86 and x86_64 for
example.
Our tests show that about 60-70% of physical memory can be allocated on a
desktop after a few days uptime. In benchmarks and stress tests, we are
finding that 80% of memory is available as contiguous blocks at the end of the
test. To compare, a standard kernel was getting < 1% of memory as large pages
on a desktop and about 8-12% of memory as large pages at the end of stress
tests.
Following this email are 12 patches that implement thie page grouping feature.
The first patch introduces a mechanism for storing flags related to a whole
block of pages. Then allocations are split between movable and all other
allocations. Following that are patches to deal with per-cpu pages and make
the mechanism configurable. The next patch moves free pages between lists
when partially allocated blocks are used for pages of another migrate type.
The second last patch groups reclaimable kernel allocations such as inode
caches together. The final patch related to groupings keeps high-order atomic
allocations.
The last two patches are more concerned with control of fragmentation. The
second last patch biases placement of non-movable allocations towards the
start of memory. This is with a view of supporting memory hot-remove of DIMMs
with higher PFNs in the future. The biasing could be enforced a lot heavier
but it would cost. The last patch agressively clusters reclaimable pages like
inode caches together.
The fragmentation reduction strategy needs to track if pages within a block
can be moved or reclaimed so that pages are freed to the appropriate list.
This patch adds a bitmap for flags affecting a whole a MAX_ORDER block of
pages.
In non-SPARSEMEM configurations, the bitmap is stored in the struct zone and
allocated during initialisation. SPARSEMEM statically allocates the bitmap in
a struct mem_section so that bitmaps do not have to be resized during memory
hotadd. This wastes a small amount of memory per unused section (usually
sizeof(unsigned long)) but the complexity of dynamically allocating the memory
is quite high.
Additional credit to Andy Whitcroft who reviewed up an earlier implementation
of the mechanism an suggested how to make it a *lot* cleaner.
Signed-off-by: Mel Gorman <mel@csn.ul.ie>
Cc: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:25:47 +04:00
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#if (MAX_ORDER - 1 + PAGE_SHIFT) > SECTION_SIZE_BITS
|
|
|
|
#error Allocator MAX_ORDER exceeds SECTION_SIZE
|
|
|
|
#endif
|
|
|
|
|
2011-05-25 04:12:33 +04:00
|
|
|
#define pfn_to_section_nr(pfn) ((pfn) >> PFN_SECTION_SHIFT)
|
|
|
|
#define section_nr_to_pfn(sec) ((sec) << PFN_SECTION_SHIFT)
|
|
|
|
|
2011-05-25 04:12:51 +04:00
|
|
|
#define SECTION_ALIGN_UP(pfn) (((pfn) + PAGES_PER_SECTION - 1) & PAGE_SECTION_MASK)
|
|
|
|
#define SECTION_ALIGN_DOWN(pfn) ((pfn) & PAGE_SECTION_MASK)
|
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
struct page;
|
2008-10-19 07:28:16 +04:00
|
|
|
struct page_cgroup;
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
struct mem_section {
|
2005-06-23 11:08:00 +04:00
|
|
|
/*
|
|
|
|
* This is, logically, a pointer to an array of struct
|
|
|
|
* pages. However, it is stored with some other magic.
|
|
|
|
* (see sparse.c::sparse_init_one_section())
|
|
|
|
*
|
2006-06-23 13:03:41 +04:00
|
|
|
* Additionally during early boot we encode node id of
|
|
|
|
* the location of the section here to guide allocation.
|
|
|
|
* (see sparse.c::memory_present())
|
|
|
|
*
|
2005-06-23 11:08:00 +04:00
|
|
|
* Making it a UL at least makes someone do a cast
|
|
|
|
* before using it wrong.
|
|
|
|
*/
|
|
|
|
unsigned long section_mem_map;
|
2007-10-16 12:25:56 +04:00
|
|
|
|
|
|
|
/* See declaration of similar field in struct zone */
|
|
|
|
unsigned long *pageblock_flags;
|
2012-08-01 03:43:02 +04:00
|
|
|
#ifdef CONFIG_MEMCG
|
2008-10-19 07:28:16 +04:00
|
|
|
/*
|
|
|
|
* If !SPARSEMEM, pgdat doesn't have page_cgroup pointer. We use
|
|
|
|
* section. (see memcontrol.h/page_cgroup.h about this.)
|
|
|
|
*/
|
|
|
|
struct page_cgroup *page_cgroup;
|
|
|
|
unsigned long pad;
|
|
|
|
#endif
|
2013-07-04 02:04:44 +04:00
|
|
|
/*
|
|
|
|
* WARNING: mem_section must be a power-of-2 in size for the
|
|
|
|
* calculation and use of SECTION_ROOT_MASK to make sense.
|
|
|
|
*/
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
};
|
|
|
|
|
2005-09-04 02:54:28 +04:00
|
|
|
#ifdef CONFIG_SPARSEMEM_EXTREME
|
|
|
|
#define SECTIONS_PER_ROOT (PAGE_SIZE / sizeof (struct mem_section))
|
|
|
|
#else
|
|
|
|
#define SECTIONS_PER_ROOT 1
|
|
|
|
#endif
|
2005-09-04 02:54:26 +04:00
|
|
|
|
2005-09-04 02:54:28 +04:00
|
|
|
#define SECTION_NR_TO_ROOT(sec) ((sec) / SECTIONS_PER_ROOT)
|
2010-05-25 01:32:47 +04:00
|
|
|
#define NR_SECTION_ROOTS DIV_ROUND_UP(NR_MEM_SECTIONS, SECTIONS_PER_ROOT)
|
2005-09-04 02:54:28 +04:00
|
|
|
#define SECTION_ROOT_MASK (SECTIONS_PER_ROOT - 1)
|
2005-09-04 02:54:26 +04:00
|
|
|
|
2005-09-04 02:54:28 +04:00
|
|
|
#ifdef CONFIG_SPARSEMEM_EXTREME
|
|
|
|
extern struct mem_section *mem_section[NR_SECTION_ROOTS];
|
2005-09-04 02:54:26 +04:00
|
|
|
#else
|
2005-09-04 02:54:28 +04:00
|
|
|
extern struct mem_section mem_section[NR_SECTION_ROOTS][SECTIONS_PER_ROOT];
|
|
|
|
#endif
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
|
2005-06-23 11:08:00 +04:00
|
|
|
static inline struct mem_section *__nr_to_section(unsigned long nr)
|
|
|
|
{
|
2005-09-04 02:54:28 +04:00
|
|
|
if (!mem_section[SECTION_NR_TO_ROOT(nr)])
|
|
|
|
return NULL;
|
|
|
|
return &mem_section[SECTION_NR_TO_ROOT(nr)][nr & SECTION_ROOT_MASK];
|
2005-06-23 11:08:00 +04:00
|
|
|
}
|
2005-10-30 04:16:51 +03:00
|
|
|
extern int __section_nr(struct mem_section* ms);
|
memory hotplug: register section/node id to free
This patch set is to free pages which is allocated by bootmem for
memory-hotremove. Some structures of memory management are allocated by
bootmem. ex) memmap, etc.
To remove memory physically, some of them must be freed according to
circumstance. This patch set makes basis to free those pages, and free
memmaps.
Basic my idea is using remain members of struct page to remember information
of users of bootmem (section number or node id). When the section is
removing, kernel can confirm it. By this information, some issues can be
solved.
1) When the memmap of removing section is allocated on other
section by bootmem, it should/can be free.
2) When the memmap of removing section is allocated on the
same section, it shouldn't be freed. Because the section has to be
logical memory offlined already and all pages must be isolated against
page allocater. If it is freed, page allocator may use it which will
be removed physically soon.
3) When removing section has other section's memmap,
kernel will be able to show easily which section should be removed
before it for user. (Not implemented yet)
4) When the above case 2), the page isolation will be able to check and skip
memmap's page when logical memory offline (offline_pages()).
Current page isolation code fails in this case because this page is
just reserved page and it can't distinguish this pages can be
removed or not. But, it will be able to do by this patch.
(Not implemented yet.)
5) The node information like pgdat has similar issues. But, this
will be able to be solved too by this.
(Not implemented yet, but, remembering node id in the pages.)
Fortunately, current bootmem allocator just keeps PageReserved flags,
and doesn't use any other members of page struct. The users of
bootmem doesn't use them too.
This patch:
This is to register information which is node or section's id. Kernel can
distinguish which node/section uses the pages allcated by bootmem. This is
basis for hot-remove sections or nodes.
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Cc: Yinghai Lu <yhlu.kernel@gmail.com>
Cc: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-04-28 13:13:31 +04:00
|
|
|
extern unsigned long usemap_size(void);
|
2005-06-23 11:08:00 +04:00
|
|
|
|
|
|
|
/*
|
|
|
|
* We use the lower bits of the mem_map pointer to store
|
|
|
|
* a little bit of information. There should be at least
|
|
|
|
* 3 bits here due to 32-bit alignment.
|
|
|
|
*/
|
|
|
|
#define SECTION_MARKED_PRESENT (1UL<<0)
|
|
|
|
#define SECTION_HAS_MEM_MAP (1UL<<1)
|
|
|
|
#define SECTION_MAP_LAST_BIT (1UL<<2)
|
|
|
|
#define SECTION_MAP_MASK (~(SECTION_MAP_LAST_BIT-1))
|
2006-06-23 13:03:41 +04:00
|
|
|
#define SECTION_NID_SHIFT 2
|
2005-06-23 11:08:00 +04:00
|
|
|
|
|
|
|
static inline struct page *__section_mem_map_addr(struct mem_section *section)
|
|
|
|
{
|
|
|
|
unsigned long map = section->section_mem_map;
|
|
|
|
map &= SECTION_MAP_MASK;
|
|
|
|
return (struct page *)map;
|
|
|
|
}
|
|
|
|
|
2007-10-16 12:24:11 +04:00
|
|
|
static inline int present_section(struct mem_section *section)
|
2005-06-23 11:08:00 +04:00
|
|
|
{
|
2005-09-04 02:54:26 +04:00
|
|
|
return (section && (section->section_mem_map & SECTION_MARKED_PRESENT));
|
2005-06-23 11:08:00 +04:00
|
|
|
}
|
|
|
|
|
2007-10-16 12:24:11 +04:00
|
|
|
static inline int present_section_nr(unsigned long nr)
|
|
|
|
{
|
|
|
|
return present_section(__nr_to_section(nr));
|
|
|
|
}
|
|
|
|
|
|
|
|
static inline int valid_section(struct mem_section *section)
|
2005-06-23 11:08:00 +04:00
|
|
|
{
|
2005-09-04 02:54:26 +04:00
|
|
|
return (section && (section->section_mem_map & SECTION_HAS_MEM_MAP));
|
2005-06-23 11:08:00 +04:00
|
|
|
}
|
|
|
|
|
|
|
|
static inline int valid_section_nr(unsigned long nr)
|
|
|
|
{
|
|
|
|
return valid_section(__nr_to_section(nr));
|
|
|
|
}
|
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
static inline struct mem_section *__pfn_to_section(unsigned long pfn)
|
|
|
|
{
|
2005-06-23 11:08:00 +04:00
|
|
|
return __nr_to_section(pfn_to_section_nr(pfn));
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
}
|
|
|
|
|
ARM: 6913/1: sparsemem: allow pfn_valid to be overridden when using SPARSEMEM
In commit eb33575c ("[ARM] Double check memmap is actually valid with a
memmap has unexpected holes V2"), a new function, memmap_valid_within,
was introduced to mmzone.h so that holes in the memmap which pass
pfn_valid in SPARSEMEM configurations can be detected and avoided.
The fix to this problem checks that the pfn <-> page linkages are
correct by calculating the page for the pfn and then checking that
page_to_pfn on that page returns the original pfn. Unfortunately, in
SPARSEMEM configurations, this results in reading from the page flags to
determine the correct section. Since the memmap here has been freed,
junk is read from memory and the check is no longer robust.
In the best case, reading from /proc/pagetypeinfo will give you the
wrong answer. In the worst case, you get SEGVs, Kernel OOPses and hung
CPUs. Furthermore, ioremap implementations that use pfn_valid to
disallow the remapping of normal memory will break.
This patch allows architectures to provide their own pfn_valid function
instead of using the default implementation used by sparsemem. The
architecture-specific version is aware of the memmap state and will
return false when passed a pfn for a freed page within a valid section.
Acked-by: Mel Gorman <mgorman@suse.de>
Acked-by: Catalin Marinas <catalin.marinas@arm.com>
Tested-by: H Hartley Sweeten <hsweeten@visionengravers.com>
Signed-off-by: Will Deacon <will.deacon@arm.com>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
2011-05-19 16:21:14 +04:00
|
|
|
#ifndef CONFIG_HAVE_ARCH_PFN_VALID
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
static inline int pfn_valid(unsigned long pfn)
|
|
|
|
{
|
|
|
|
if (pfn_to_section_nr(pfn) >= NR_MEM_SECTIONS)
|
|
|
|
return 0;
|
2005-06-23 11:08:00 +04:00
|
|
|
return valid_section(__nr_to_section(pfn_to_section_nr(pfn)));
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
}
|
ARM: 6913/1: sparsemem: allow pfn_valid to be overridden when using SPARSEMEM
In commit eb33575c ("[ARM] Double check memmap is actually valid with a
memmap has unexpected holes V2"), a new function, memmap_valid_within,
was introduced to mmzone.h so that holes in the memmap which pass
pfn_valid in SPARSEMEM configurations can be detected and avoided.
The fix to this problem checks that the pfn <-> page linkages are
correct by calculating the page for the pfn and then checking that
page_to_pfn on that page returns the original pfn. Unfortunately, in
SPARSEMEM configurations, this results in reading from the page flags to
determine the correct section. Since the memmap here has been freed,
junk is read from memory and the check is no longer robust.
In the best case, reading from /proc/pagetypeinfo will give you the
wrong answer. In the worst case, you get SEGVs, Kernel OOPses and hung
CPUs. Furthermore, ioremap implementations that use pfn_valid to
disallow the remapping of normal memory will break.
This patch allows architectures to provide their own pfn_valid function
instead of using the default implementation used by sparsemem. The
architecture-specific version is aware of the memmap state and will
return false when passed a pfn for a freed page within a valid section.
Acked-by: Mel Gorman <mgorman@suse.de>
Acked-by: Catalin Marinas <catalin.marinas@arm.com>
Tested-by: H Hartley Sweeten <hsweeten@visionengravers.com>
Signed-off-by: Will Deacon <will.deacon@arm.com>
Signed-off-by: Russell King <rmk+kernel@arm.linux.org.uk>
2011-05-19 16:21:14 +04:00
|
|
|
#endif
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
|
2007-10-16 12:24:11 +04:00
|
|
|
static inline int pfn_present(unsigned long pfn)
|
|
|
|
{
|
|
|
|
if (pfn_to_section_nr(pfn) >= NR_MEM_SECTIONS)
|
|
|
|
return 0;
|
|
|
|
return present_section(__nr_to_section(pfn_to_section_nr(pfn)));
|
|
|
|
}
|
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
/*
|
|
|
|
* These are _only_ used during initialisation, therefore they
|
|
|
|
* can use __initdata ... They could have names to indicate
|
|
|
|
* this restriction.
|
|
|
|
*/
|
|
|
|
#ifdef CONFIG_NUMA
|
2006-01-06 11:10:54 +03:00
|
|
|
#define pfn_to_nid(pfn) \
|
|
|
|
({ \
|
|
|
|
unsigned long __pfn_to_nid_pfn = (pfn); \
|
|
|
|
page_to_nid(pfn_to_page(__pfn_to_nid_pfn)); \
|
|
|
|
})
|
2006-01-06 11:10:53 +03:00
|
|
|
#else
|
|
|
|
#define pfn_to_nid(pfn) (0)
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#endif
|
|
|
|
|
|
|
|
#define early_pfn_valid(pfn) pfn_valid(pfn)
|
|
|
|
void sparse_init(void);
|
|
|
|
#else
|
|
|
|
#define sparse_init() do {} while (0)
|
2005-09-04 02:54:29 +04:00
|
|
|
#define sparse_index_init(_sec, _nid) do {} while (0)
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#endif /* CONFIG_SPARSEMEM */
|
|
|
|
|
2006-10-21 21:24:14 +04:00
|
|
|
#ifdef CONFIG_NODES_SPAN_OTHER_NODES
|
2009-02-19 01:48:33 +03:00
|
|
|
bool early_pfn_in_nid(unsigned long pfn, int nid);
|
2006-10-21 21:24:14 +04:00
|
|
|
#else
|
|
|
|
#define early_pfn_in_nid(pfn, nid) (1)
|
|
|
|
#endif
|
|
|
|
|
[PATCH] sparsemem memory model
Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of
mem_map[] is needed by discontiguous memory machines (like in the old
CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem
replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually
become a complete replacement.
A significant advantage over DISCONTIGMEM is that it's completely separated
from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA
and DISCONTIG are often confused.
Another advantage is that sparse doesn't require each NUMA node's ranges to be
contiguous. It can handle overlapping ranges between nodes with no problems,
where DISCONTIGMEM currently throws away that memory.
Sparsemem uses an array to provide different pfn_to_page() translations for
each SECTION_SIZE area of physical memory. This is what allows the mem_map[]
to be chopped up.
In order to do quick pfn_to_page() operations, the section number of the page
is encoded in page->flags. Part of the sparsemem infrastructure enables
sharing of these bits more dynamically (at compile-time) between the
page_zone() and sparsemem operations. However, on 32-bit architectures, the
number of bits is quite limited, and may require growing the size of the
page->flags type in certain conditions. Several things might force this to
occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of
memory), an increase in the physical address space, or an increase in the
number of used page->flags.
One thing to note is that, once sparsemem is present, the NUMA node
information no longer needs to be stored in the page->flags. It might provide
speed increases on certain platforms and will be stored there if there is
room. But, if out of room, an alternate (theoretically slower) mechanism is
used.
This patch introduces CONFIG_FLATMEM. It is used in almost all cases where
there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM
often have to compile out the same areas of code.
Signed-off-by: Andy Whitcroft <apw@shadowen.org>
Signed-off-by: Dave Hansen <haveblue@us.ibm.com>
Signed-off-by: Martin Bligh <mbligh@aracnet.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: Bob Picco <bob.picco@hp.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 11:07:54 +04:00
|
|
|
#ifndef early_pfn_valid
|
|
|
|
#define early_pfn_valid(pfn) (1)
|
|
|
|
#endif
|
|
|
|
|
|
|
|
void memory_present(int nid, unsigned long start, unsigned long end);
|
|
|
|
unsigned long __init node_memmap_size_bytes(int, unsigned long, unsigned long);
|
|
|
|
|
2007-05-07 01:49:14 +04:00
|
|
|
/*
|
|
|
|
* If it is possible to have holes within a MAX_ORDER_NR_PAGES, then we
|
|
|
|
* need to check pfn validility within that MAX_ORDER_NR_PAGES block.
|
|
|
|
* pfn_valid_within() should be used in this case; we optimise this away
|
|
|
|
* when we have no holes within a MAX_ORDER_NR_PAGES block.
|
|
|
|
*/
|
|
|
|
#ifdef CONFIG_HOLES_IN_ZONE
|
|
|
|
#define pfn_valid_within(pfn) pfn_valid(pfn)
|
|
|
|
#else
|
|
|
|
#define pfn_valid_within(pfn) (1)
|
|
|
|
#endif
|
|
|
|
|
2009-05-13 20:34:48 +04:00
|
|
|
#ifdef CONFIG_ARCH_HAS_HOLES_MEMORYMODEL
|
|
|
|
/*
|
|
|
|
* pfn_valid() is meant to be able to tell if a given PFN has valid memmap
|
|
|
|
* associated with it or not. In FLATMEM, it is expected that holes always
|
|
|
|
* have valid memmap as long as there is valid PFNs either side of the hole.
|
|
|
|
* In SPARSEMEM, it is assumed that a valid section has a memmap for the
|
|
|
|
* entire section.
|
|
|
|
*
|
|
|
|
* However, an ARM, and maybe other embedded architectures in the future
|
|
|
|
* free memmap backing holes to save memory on the assumption the memmap is
|
|
|
|
* never used. The page_zone linkages are then broken even though pfn_valid()
|
|
|
|
* returns true. A walker of the full memmap must then do this additional
|
|
|
|
* check to ensure the memmap they are looking at is sane by making sure
|
|
|
|
* the zone and PFN linkages are still valid. This is expensive, but walkers
|
|
|
|
* of the full memmap are extremely rare.
|
|
|
|
*/
|
|
|
|
int memmap_valid_within(unsigned long pfn,
|
|
|
|
struct page *page, struct zone *zone);
|
|
|
|
#else
|
|
|
|
static inline int memmap_valid_within(unsigned long pfn,
|
|
|
|
struct page *page, struct zone *zone)
|
|
|
|
{
|
|
|
|
return 1;
|
|
|
|
}
|
|
|
|
#endif /* CONFIG_ARCH_HAS_HOLES_MEMORYMODEL */
|
|
|
|
|
2008-04-28 13:12:54 +04:00
|
|
|
#endif /* !__GENERATING_BOUNDS.H */
|
2005-04-17 02:20:36 +04:00
|
|
|
#endif /* !__ASSEMBLY__ */
|
|
|
|
#endif /* _LINUX_MMZONE_H */
|