WSL2-Linux-Kernel/include/linux/slub_def.h

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License cleanup: add SPDX GPL-2.0 license identifier to files with no license Many source files in the tree are missing licensing information, which makes it harder for compliance tools to determine the correct license. By default all files without license information are under the default license of the kernel, which is GPL version 2. Update the files which contain no license information with the 'GPL-2.0' SPDX license identifier. The SPDX identifier is a legally binding shorthand, which can be used instead of the full boiler plate text. This patch is based on work done by Thomas Gleixner and Kate Stewart and Philippe Ombredanne. How this work was done: Patches were generated and checked against linux-4.14-rc6 for a subset of the use cases: - file had no licensing information it it. - file was a */uapi/* one with no licensing information in it, - file was a */uapi/* one with existing licensing information, Further patches will be generated in subsequent months to fix up cases where non-standard license headers were used, and references to license had to be inferred by heuristics based on keywords. The analysis to determine which SPDX License Identifier to be applied to a file was done in a spreadsheet of side by side results from of the output of two independent scanners (ScanCode & Windriver) producing SPDX tag:value files created by Philippe Ombredanne. Philippe prepared the base worksheet, and did an initial spot review of a few 1000 files. The 4.13 kernel was the starting point of the analysis with 60,537 files assessed. Kate Stewart did a file by file comparison of the scanner results in the spreadsheet to determine which SPDX license identifier(s) to be applied to the file. She confirmed any determination that was not immediately clear with lawyers working with the Linux Foundation. Criteria used to select files for SPDX license identifier tagging was: - Files considered eligible had to be source code files. - Make and config files were included as candidates if they contained >5 lines of source - File already had some variant of a license header in it (even if <5 lines). All documentation files were explicitly excluded. The following heuristics were used to determine which SPDX license identifiers to apply. - when both scanners couldn't find any license traces, file was considered to have no license information in it, and the top level COPYING file license applied. For non */uapi/* files that summary was: SPDX license identifier # files ---------------------------------------------------|------- GPL-2.0 11139 and resulted in the first patch in this series. If that file was a */uapi/* path one, it was "GPL-2.0 WITH Linux-syscall-note" otherwise it was "GPL-2.0". Results of that was: SPDX license identifier # files ---------------------------------------------------|------- GPL-2.0 WITH Linux-syscall-note 930 and resulted in the second patch in this series. - if a file had some form of licensing information in it, and was one of the */uapi/* ones, it was denoted with the Linux-syscall-note if any GPL family license was found in the file or had no licensing in it (per prior point). Results summary: SPDX license identifier # files ---------------------------------------------------|------ GPL-2.0 WITH Linux-syscall-note 270 GPL-2.0+ WITH Linux-syscall-note 169 ((GPL-2.0 WITH Linux-syscall-note) OR BSD-2-Clause) 21 ((GPL-2.0 WITH Linux-syscall-note) OR BSD-3-Clause) 17 LGPL-2.1+ WITH Linux-syscall-note 15 GPL-1.0+ WITH Linux-syscall-note 14 ((GPL-2.0+ WITH Linux-syscall-note) OR BSD-3-Clause) 5 LGPL-2.0+ WITH Linux-syscall-note 4 LGPL-2.1 WITH Linux-syscall-note 3 ((GPL-2.0 WITH Linux-syscall-note) OR MIT) 3 ((GPL-2.0 WITH Linux-syscall-note) AND MIT) 1 and that resulted in the third patch in this series. - when the two scanners agreed on the detected license(s), that became the concluded license(s). - when there was disagreement between the two scanners (one detected a license but the other didn't, or they both detected different licenses) a manual inspection of the file occurred. - In most cases a manual inspection of the information in the file resulted in a clear resolution of the license that should apply (and which scanner probably needed to revisit its heuristics). - When it was not immediately clear, the license identifier was confirmed with lawyers working with the Linux Foundation. - If there was any question as to the appropriate license identifier, the file was flagged for further research and to be revisited later in time. In total, over 70 hours of logged manual review was done on the spreadsheet to determine the SPDX license identifiers to apply to the source files by Kate, Philippe, Thomas and, in some cases, confirmation by lawyers working with the Linux Foundation. Kate also obtained a third independent scan of the 4.13 code base from FOSSology, and compared selected files where the other two scanners disagreed against that SPDX file, to see if there was new insights. The Windriver scanner is based on an older version of FOSSology in part, so they are related. Thomas did random spot checks in about 500 files from the spreadsheets for the uapi headers and agreed with SPDX license identifier in the files he inspected. For the non-uapi files Thomas did random spot checks in about 15000 files. In initial set of patches against 4.14-rc6, 3 files were found to have copy/paste license identifier errors, and have been fixed to reflect the correct identifier. Additionally Philippe spent 10 hours this week doing a detailed manual inspection and review of the 12,461 patched files from the initial patch version early this week with: - a full scancode scan run, collecting the matched texts, detected license ids and scores - reviewing anything where there was a license detected (about 500+ files) to ensure that the applied SPDX license was correct - reviewing anything where there was no detection but the patch license was not GPL-2.0 WITH Linux-syscall-note to ensure that the applied SPDX license was correct This produced a worksheet with 20 files needing minor correction. This worksheet was then exported into 3 different .csv files for the different types of files to be modified. These .csv files were then reviewed by Greg. Thomas wrote a script to parse the csv files and add the proper SPDX tag to the file, in the format that the file expected. This script was further refined by Greg based on the output to detect more types of files automatically and to distinguish between header and source .c files (which need different comment types.) Finally Greg ran the script using the .csv files to generate the patches. Reviewed-by: Kate Stewart <kstewart@linuxfoundation.org> Reviewed-by: Philippe Ombredanne <pombredanne@nexb.com> Reviewed-by: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: Greg Kroah-Hartman <gregkh@linuxfoundation.org>
2017-11-01 17:07:57 +03:00
/* SPDX-License-Identifier: GPL-2.0 */
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
#ifndef _LINUX_SLUB_DEF_H
#define _LINUX_SLUB_DEF_H
/*
* SLUB : A Slab allocator without object queues.
*
* (C) 2007 SGI, Christoph Lameter
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
*/
#include <linux/kobject.h>
SLUB: Support for performance statistics The statistics provided here allow the monitoring of allocator behavior but at the cost of some (minimal) loss of performance. Counters are placed in SLUB's per cpu data structure. The per cpu structure may be extended by the statistics to grow larger than one cacheline which will increase the cache footprint of SLUB. There is a compile option to enable/disable the inclusion of the runtime statistics and its off by default. The slabinfo tool is enhanced to support these statistics via two options: -D Switches the line of information displayed for a slab from size mode to activity mode. -A Sorts the slabs displayed by activity. This allows the display of the slabs most important to the performance of a certain load. -r Report option will report detailed statistics on Example (tbench load): slabinfo -AD ->Shows the most active slabs Name Objects Alloc Free %Fast skbuff_fclone_cache 33 111953835 111953835 99 99 :0000192 2666 5283688 5281047 99 99 :0001024 849 5247230 5246389 83 83 vm_area_struct 1349 119642 118355 91 22 :0004096 15 66753 66751 98 98 :0000064 2067 25297 23383 98 78 dentry 10259 28635 18464 91 45 :0000080 11004 18950 8089 98 98 :0000096 1703 12358 10784 99 98 :0000128 762 10582 9875 94 18 :0000512 184 9807 9647 95 81 :0002048 479 9669 9195 83 65 anon_vma 777 9461 9002 99 71 kmalloc-8 6492 9981 5624 99 97 :0000768 258 7174 6931 58 15 So the skbuff_fclone_cache is of highest importance for the tbench load. Pretty high load on the 192 sized slab. Look for the aliases slabinfo -a | grep 000192 :0000192 <- xfs_btree_cur filp kmalloc-192 uid_cache tw_sock_TCP request_sock_TCPv6 tw_sock_TCPv6 skbuff_head_cache xfs_ili Likely skbuff_head_cache. Looking into the statistics of the skbuff_fclone_cache is possible through slabinfo skbuff_fclone_cache ->-r option implied if cache name is mentioned .... Usual output ... Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 111953360 111946981 99 99 Slowpath 1044 7423 0 0 Page Alloc 272 264 0 0 Add partial 25 325 0 0 Remove partial 86 264 0 0 RemoteObj/SlabFrozen 350 4832 0 0 Total 111954404 111954404 Flushes 49 Refill 0 Deactivate Full=325(92%) Empty=0(0%) ToHead=24(6%) ToTail=1(0%) Looks good because the fastpath is overwhelmingly taken. skbuff_head_cache: Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 5297262 5259882 99 99 Slowpath 4477 39586 0 0 Page Alloc 937 824 0 0 Add partial 0 2515 0 0 Remove partial 1691 824 0 0 RemoteObj/SlabFrozen 2621 9684 0 0 Total 5301739 5299468 Deactivate Full=2620(100%) Empty=0(0%) ToHead=0(0%) ToTail=0(0%) Descriptions of the output: Total: The total number of allocation and frees that occurred for a slab Fastpath: The number of allocations/frees that used the fastpath. Slowpath: Other allocations Page Alloc: Number of calls to the page allocator as a result of slowpath processing Add Partial: Number of slabs added to the partial list through free or alloc (occurs during cpuslab flushes) Remove Partial: Number of slabs removed from the partial list as a result of allocations retrieving a partial slab or by a free freeing the last object of a slab. RemoteObj/Froz: How many times were remotely freed object encountered when a slab was about to be deactivated. Frozen: How many times was free able to skip list processing because the slab was in use as the cpuslab of another processor. Flushes: Number of times the cpuslab was flushed on request (kmem_cache_shrink, may result from races in __slab_alloc) Refill: Number of times we were able to refill the cpuslab from remotely freed objects for the same slab. Deactivate: Statistics how slabs were deactivated. Shows how they were put onto the partial list. In general fastpath is very good. Slowpath without partial list processing is also desirable. Any touching of partial list uses node specific locks which may potentially cause list lock contention. Signed-off-by: Christoph Lameter <clameter@sgi.com>
2008-02-08 04:47:41 +03:00
enum stat_item {
ALLOC_FASTPATH, /* Allocation from cpu slab */
ALLOC_SLOWPATH, /* Allocation by getting a new cpu slab */
FREE_FASTPATH, /* Free to cpu slab */
SLUB: Support for performance statistics The statistics provided here allow the monitoring of allocator behavior but at the cost of some (minimal) loss of performance. Counters are placed in SLUB's per cpu data structure. The per cpu structure may be extended by the statistics to grow larger than one cacheline which will increase the cache footprint of SLUB. There is a compile option to enable/disable the inclusion of the runtime statistics and its off by default. The slabinfo tool is enhanced to support these statistics via two options: -D Switches the line of information displayed for a slab from size mode to activity mode. -A Sorts the slabs displayed by activity. This allows the display of the slabs most important to the performance of a certain load. -r Report option will report detailed statistics on Example (tbench load): slabinfo -AD ->Shows the most active slabs Name Objects Alloc Free %Fast skbuff_fclone_cache 33 111953835 111953835 99 99 :0000192 2666 5283688 5281047 99 99 :0001024 849 5247230 5246389 83 83 vm_area_struct 1349 119642 118355 91 22 :0004096 15 66753 66751 98 98 :0000064 2067 25297 23383 98 78 dentry 10259 28635 18464 91 45 :0000080 11004 18950 8089 98 98 :0000096 1703 12358 10784 99 98 :0000128 762 10582 9875 94 18 :0000512 184 9807 9647 95 81 :0002048 479 9669 9195 83 65 anon_vma 777 9461 9002 99 71 kmalloc-8 6492 9981 5624 99 97 :0000768 258 7174 6931 58 15 So the skbuff_fclone_cache is of highest importance for the tbench load. Pretty high load on the 192 sized slab. Look for the aliases slabinfo -a | grep 000192 :0000192 <- xfs_btree_cur filp kmalloc-192 uid_cache tw_sock_TCP request_sock_TCPv6 tw_sock_TCPv6 skbuff_head_cache xfs_ili Likely skbuff_head_cache. Looking into the statistics of the skbuff_fclone_cache is possible through slabinfo skbuff_fclone_cache ->-r option implied if cache name is mentioned .... Usual output ... Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 111953360 111946981 99 99 Slowpath 1044 7423 0 0 Page Alloc 272 264 0 0 Add partial 25 325 0 0 Remove partial 86 264 0 0 RemoteObj/SlabFrozen 350 4832 0 0 Total 111954404 111954404 Flushes 49 Refill 0 Deactivate Full=325(92%) Empty=0(0%) ToHead=24(6%) ToTail=1(0%) Looks good because the fastpath is overwhelmingly taken. skbuff_head_cache: Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 5297262 5259882 99 99 Slowpath 4477 39586 0 0 Page Alloc 937 824 0 0 Add partial 0 2515 0 0 Remove partial 1691 824 0 0 RemoteObj/SlabFrozen 2621 9684 0 0 Total 5301739 5299468 Deactivate Full=2620(100%) Empty=0(0%) ToHead=0(0%) ToTail=0(0%) Descriptions of the output: Total: The total number of allocation and frees that occurred for a slab Fastpath: The number of allocations/frees that used the fastpath. Slowpath: Other allocations Page Alloc: Number of calls to the page allocator as a result of slowpath processing Add Partial: Number of slabs added to the partial list through free or alloc (occurs during cpuslab flushes) Remove Partial: Number of slabs removed from the partial list as a result of allocations retrieving a partial slab or by a free freeing the last object of a slab. RemoteObj/Froz: How many times were remotely freed object encountered when a slab was about to be deactivated. Frozen: How many times was free able to skip list processing because the slab was in use as the cpuslab of another processor. Flushes: Number of times the cpuslab was flushed on request (kmem_cache_shrink, may result from races in __slab_alloc) Refill: Number of times we were able to refill the cpuslab from remotely freed objects for the same slab. Deactivate: Statistics how slabs were deactivated. Shows how they were put onto the partial list. In general fastpath is very good. Slowpath without partial list processing is also desirable. Any touching of partial list uses node specific locks which may potentially cause list lock contention. Signed-off-by: Christoph Lameter <clameter@sgi.com>
2008-02-08 04:47:41 +03:00
FREE_SLOWPATH, /* Freeing not to cpu slab */
FREE_FROZEN, /* Freeing to frozen slab */
FREE_ADD_PARTIAL, /* Freeing moves slab to partial list */
FREE_REMOVE_PARTIAL, /* Freeing removes last object */
ALLOC_FROM_PARTIAL, /* Cpu slab acquired from node partial list */
SLUB: Support for performance statistics The statistics provided here allow the monitoring of allocator behavior but at the cost of some (minimal) loss of performance. Counters are placed in SLUB's per cpu data structure. The per cpu structure may be extended by the statistics to grow larger than one cacheline which will increase the cache footprint of SLUB. There is a compile option to enable/disable the inclusion of the runtime statistics and its off by default. The slabinfo tool is enhanced to support these statistics via two options: -D Switches the line of information displayed for a slab from size mode to activity mode. -A Sorts the slabs displayed by activity. This allows the display of the slabs most important to the performance of a certain load. -r Report option will report detailed statistics on Example (tbench load): slabinfo -AD ->Shows the most active slabs Name Objects Alloc Free %Fast skbuff_fclone_cache 33 111953835 111953835 99 99 :0000192 2666 5283688 5281047 99 99 :0001024 849 5247230 5246389 83 83 vm_area_struct 1349 119642 118355 91 22 :0004096 15 66753 66751 98 98 :0000064 2067 25297 23383 98 78 dentry 10259 28635 18464 91 45 :0000080 11004 18950 8089 98 98 :0000096 1703 12358 10784 99 98 :0000128 762 10582 9875 94 18 :0000512 184 9807 9647 95 81 :0002048 479 9669 9195 83 65 anon_vma 777 9461 9002 99 71 kmalloc-8 6492 9981 5624 99 97 :0000768 258 7174 6931 58 15 So the skbuff_fclone_cache is of highest importance for the tbench load. Pretty high load on the 192 sized slab. Look for the aliases slabinfo -a | grep 000192 :0000192 <- xfs_btree_cur filp kmalloc-192 uid_cache tw_sock_TCP request_sock_TCPv6 tw_sock_TCPv6 skbuff_head_cache xfs_ili Likely skbuff_head_cache. Looking into the statistics of the skbuff_fclone_cache is possible through slabinfo skbuff_fclone_cache ->-r option implied if cache name is mentioned .... Usual output ... Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 111953360 111946981 99 99 Slowpath 1044 7423 0 0 Page Alloc 272 264 0 0 Add partial 25 325 0 0 Remove partial 86 264 0 0 RemoteObj/SlabFrozen 350 4832 0 0 Total 111954404 111954404 Flushes 49 Refill 0 Deactivate Full=325(92%) Empty=0(0%) ToHead=24(6%) ToTail=1(0%) Looks good because the fastpath is overwhelmingly taken. skbuff_head_cache: Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 5297262 5259882 99 99 Slowpath 4477 39586 0 0 Page Alloc 937 824 0 0 Add partial 0 2515 0 0 Remove partial 1691 824 0 0 RemoteObj/SlabFrozen 2621 9684 0 0 Total 5301739 5299468 Deactivate Full=2620(100%) Empty=0(0%) ToHead=0(0%) ToTail=0(0%) Descriptions of the output: Total: The total number of allocation and frees that occurred for a slab Fastpath: The number of allocations/frees that used the fastpath. Slowpath: Other allocations Page Alloc: Number of calls to the page allocator as a result of slowpath processing Add Partial: Number of slabs added to the partial list through free or alloc (occurs during cpuslab flushes) Remove Partial: Number of slabs removed from the partial list as a result of allocations retrieving a partial slab or by a free freeing the last object of a slab. RemoteObj/Froz: How many times were remotely freed object encountered when a slab was about to be deactivated. Frozen: How many times was free able to skip list processing because the slab was in use as the cpuslab of another processor. Flushes: Number of times the cpuslab was flushed on request (kmem_cache_shrink, may result from races in __slab_alloc) Refill: Number of times we were able to refill the cpuslab from remotely freed objects for the same slab. Deactivate: Statistics how slabs were deactivated. Shows how they were put onto the partial list. In general fastpath is very good. Slowpath without partial list processing is also desirable. Any touching of partial list uses node specific locks which may potentially cause list lock contention. Signed-off-by: Christoph Lameter <clameter@sgi.com>
2008-02-08 04:47:41 +03:00
ALLOC_SLAB, /* Cpu slab acquired from page allocator */
ALLOC_REFILL, /* Refill cpu slab from slab freelist */
ALLOC_NODE_MISMATCH, /* Switching cpu slab */
SLUB: Support for performance statistics The statistics provided here allow the monitoring of allocator behavior but at the cost of some (minimal) loss of performance. Counters are placed in SLUB's per cpu data structure. The per cpu structure may be extended by the statistics to grow larger than one cacheline which will increase the cache footprint of SLUB. There is a compile option to enable/disable the inclusion of the runtime statistics and its off by default. The slabinfo tool is enhanced to support these statistics via two options: -D Switches the line of information displayed for a slab from size mode to activity mode. -A Sorts the slabs displayed by activity. This allows the display of the slabs most important to the performance of a certain load. -r Report option will report detailed statistics on Example (tbench load): slabinfo -AD ->Shows the most active slabs Name Objects Alloc Free %Fast skbuff_fclone_cache 33 111953835 111953835 99 99 :0000192 2666 5283688 5281047 99 99 :0001024 849 5247230 5246389 83 83 vm_area_struct 1349 119642 118355 91 22 :0004096 15 66753 66751 98 98 :0000064 2067 25297 23383 98 78 dentry 10259 28635 18464 91 45 :0000080 11004 18950 8089 98 98 :0000096 1703 12358 10784 99 98 :0000128 762 10582 9875 94 18 :0000512 184 9807 9647 95 81 :0002048 479 9669 9195 83 65 anon_vma 777 9461 9002 99 71 kmalloc-8 6492 9981 5624 99 97 :0000768 258 7174 6931 58 15 So the skbuff_fclone_cache is of highest importance for the tbench load. Pretty high load on the 192 sized slab. Look for the aliases slabinfo -a | grep 000192 :0000192 <- xfs_btree_cur filp kmalloc-192 uid_cache tw_sock_TCP request_sock_TCPv6 tw_sock_TCPv6 skbuff_head_cache xfs_ili Likely skbuff_head_cache. Looking into the statistics of the skbuff_fclone_cache is possible through slabinfo skbuff_fclone_cache ->-r option implied if cache name is mentioned .... Usual output ... Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 111953360 111946981 99 99 Slowpath 1044 7423 0 0 Page Alloc 272 264 0 0 Add partial 25 325 0 0 Remove partial 86 264 0 0 RemoteObj/SlabFrozen 350 4832 0 0 Total 111954404 111954404 Flushes 49 Refill 0 Deactivate Full=325(92%) Empty=0(0%) ToHead=24(6%) ToTail=1(0%) Looks good because the fastpath is overwhelmingly taken. skbuff_head_cache: Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 5297262 5259882 99 99 Slowpath 4477 39586 0 0 Page Alloc 937 824 0 0 Add partial 0 2515 0 0 Remove partial 1691 824 0 0 RemoteObj/SlabFrozen 2621 9684 0 0 Total 5301739 5299468 Deactivate Full=2620(100%) Empty=0(0%) ToHead=0(0%) ToTail=0(0%) Descriptions of the output: Total: The total number of allocation and frees that occurred for a slab Fastpath: The number of allocations/frees that used the fastpath. Slowpath: Other allocations Page Alloc: Number of calls to the page allocator as a result of slowpath processing Add Partial: Number of slabs added to the partial list through free or alloc (occurs during cpuslab flushes) Remove Partial: Number of slabs removed from the partial list as a result of allocations retrieving a partial slab or by a free freeing the last object of a slab. RemoteObj/Froz: How many times were remotely freed object encountered when a slab was about to be deactivated. Frozen: How many times was free able to skip list processing because the slab was in use as the cpuslab of another processor. Flushes: Number of times the cpuslab was flushed on request (kmem_cache_shrink, may result from races in __slab_alloc) Refill: Number of times we were able to refill the cpuslab from remotely freed objects for the same slab. Deactivate: Statistics how slabs were deactivated. Shows how they were put onto the partial list. In general fastpath is very good. Slowpath without partial list processing is also desirable. Any touching of partial list uses node specific locks which may potentially cause list lock contention. Signed-off-by: Christoph Lameter <clameter@sgi.com>
2008-02-08 04:47:41 +03:00
FREE_SLAB, /* Slab freed to the page allocator */
CPUSLAB_FLUSH, /* Abandoning of the cpu slab */
DEACTIVATE_FULL, /* Cpu slab was full when deactivated */
DEACTIVATE_EMPTY, /* Cpu slab was empty when deactivated */
DEACTIVATE_TO_HEAD, /* Cpu slab was moved to the head of partials */
DEACTIVATE_TO_TAIL, /* Cpu slab was moved to the tail of partials */
DEACTIVATE_REMOTE_FREES,/* Slab contained remotely freed objects */
DEACTIVATE_BYPASS, /* Implicit deactivation */
ORDER_FALLBACK, /* Number of times fallback was necessary */
CMPXCHG_DOUBLE_CPU_FAIL,/* Failure of this_cpu_cmpxchg_double */
CMPXCHG_DOUBLE_FAIL, /* Number of times that cmpxchg double did not match */
slub: per cpu cache for partial pages Allow filling out the rest of the kmem_cache_cpu cacheline with pointers to partial pages. The partial page list is used in slab_free() to avoid per node lock taking. In __slab_alloc() we can then take multiple partial pages off the per node partial list in one go reducing node lock pressure. We can also use the per cpu partial list in slab_alloc() to avoid scanning partial lists for pages with free objects. The main effect of a per cpu partial list is that the per node list_lock is taken for batches of partial pages instead of individual ones. Potential future enhancements: 1. The pickup from the partial list could be perhaps be done without disabling interrupts with some work. The free path already puts the page into the per cpu partial list without disabling interrupts. 2. __slab_free() may have some code paths that could use optimization. Performance: Before After ./hackbench 100 process 200000 Time: 1953.047 1564.614 ./hackbench 100 process 20000 Time: 207.176 156.940 ./hackbench 100 process 20000 Time: 204.468 156.940 ./hackbench 100 process 20000 Time: 204.879 158.772 ./hackbench 10 process 20000 Time: 20.153 15.853 ./hackbench 10 process 20000 Time: 20.153 15.986 ./hackbench 10 process 20000 Time: 19.363 16.111 ./hackbench 1 process 20000 Time: 2.518 2.307 ./hackbench 1 process 20000 Time: 2.258 2.339 ./hackbench 1 process 20000 Time: 2.864 2.163 Signed-off-by: Christoph Lameter <cl@linux.com> Signed-off-by: Pekka Enberg <penberg@kernel.org>
2011-08-10 01:12:27 +04:00
CPU_PARTIAL_ALLOC, /* Used cpu partial on alloc */
CPU_PARTIAL_FREE, /* Refill cpu partial on free */
CPU_PARTIAL_NODE, /* Refill cpu partial from node partial */
CPU_PARTIAL_DRAIN, /* Drain cpu partial to node partial */
SLUB: Support for performance statistics The statistics provided here allow the monitoring of allocator behavior but at the cost of some (minimal) loss of performance. Counters are placed in SLUB's per cpu data structure. The per cpu structure may be extended by the statistics to grow larger than one cacheline which will increase the cache footprint of SLUB. There is a compile option to enable/disable the inclusion of the runtime statistics and its off by default. The slabinfo tool is enhanced to support these statistics via two options: -D Switches the line of information displayed for a slab from size mode to activity mode. -A Sorts the slabs displayed by activity. This allows the display of the slabs most important to the performance of a certain load. -r Report option will report detailed statistics on Example (tbench load): slabinfo -AD ->Shows the most active slabs Name Objects Alloc Free %Fast skbuff_fclone_cache 33 111953835 111953835 99 99 :0000192 2666 5283688 5281047 99 99 :0001024 849 5247230 5246389 83 83 vm_area_struct 1349 119642 118355 91 22 :0004096 15 66753 66751 98 98 :0000064 2067 25297 23383 98 78 dentry 10259 28635 18464 91 45 :0000080 11004 18950 8089 98 98 :0000096 1703 12358 10784 99 98 :0000128 762 10582 9875 94 18 :0000512 184 9807 9647 95 81 :0002048 479 9669 9195 83 65 anon_vma 777 9461 9002 99 71 kmalloc-8 6492 9981 5624 99 97 :0000768 258 7174 6931 58 15 So the skbuff_fclone_cache is of highest importance for the tbench load. Pretty high load on the 192 sized slab. Look for the aliases slabinfo -a | grep 000192 :0000192 <- xfs_btree_cur filp kmalloc-192 uid_cache tw_sock_TCP request_sock_TCPv6 tw_sock_TCPv6 skbuff_head_cache xfs_ili Likely skbuff_head_cache. Looking into the statistics of the skbuff_fclone_cache is possible through slabinfo skbuff_fclone_cache ->-r option implied if cache name is mentioned .... Usual output ... Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 111953360 111946981 99 99 Slowpath 1044 7423 0 0 Page Alloc 272 264 0 0 Add partial 25 325 0 0 Remove partial 86 264 0 0 RemoteObj/SlabFrozen 350 4832 0 0 Total 111954404 111954404 Flushes 49 Refill 0 Deactivate Full=325(92%) Empty=0(0%) ToHead=24(6%) ToTail=1(0%) Looks good because the fastpath is overwhelmingly taken. skbuff_head_cache: Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 5297262 5259882 99 99 Slowpath 4477 39586 0 0 Page Alloc 937 824 0 0 Add partial 0 2515 0 0 Remove partial 1691 824 0 0 RemoteObj/SlabFrozen 2621 9684 0 0 Total 5301739 5299468 Deactivate Full=2620(100%) Empty=0(0%) ToHead=0(0%) ToTail=0(0%) Descriptions of the output: Total: The total number of allocation and frees that occurred for a slab Fastpath: The number of allocations/frees that used the fastpath. Slowpath: Other allocations Page Alloc: Number of calls to the page allocator as a result of slowpath processing Add Partial: Number of slabs added to the partial list through free or alloc (occurs during cpuslab flushes) Remove Partial: Number of slabs removed from the partial list as a result of allocations retrieving a partial slab or by a free freeing the last object of a slab. RemoteObj/Froz: How many times were remotely freed object encountered when a slab was about to be deactivated. Frozen: How many times was free able to skip list processing because the slab was in use as the cpuslab of another processor. Flushes: Number of times the cpuslab was flushed on request (kmem_cache_shrink, may result from races in __slab_alloc) Refill: Number of times we were able to refill the cpuslab from remotely freed objects for the same slab. Deactivate: Statistics how slabs were deactivated. Shows how they were put onto the partial list. In general fastpath is very good. Slowpath without partial list processing is also desirable. Any touching of partial list uses node specific locks which may potentially cause list lock contention. Signed-off-by: Christoph Lameter <clameter@sgi.com>
2008-02-08 04:47:41 +03:00
NR_SLUB_STAT_ITEMS };
SLUB: Avoid page struct cacheline bouncing due to remote frees to cpu slab A remote free may access the same page struct that also contains the lockless freelist for the cpu slab. If objects have a short lifetime and are freed by a different processor then remote frees back to the slab from which we are currently allocating are frequent. The cacheline with the page struct needs to be repeately acquired in exclusive mode by both the allocating thread and the freeing thread. If this is frequent enough then performance will suffer because of cacheline bouncing. This patchset puts the lockless_freelist pointer in its own cacheline. In order to make that happen we introduce a per cpu structure called kmem_cache_cpu. Instead of keeping an array of pointers to page structs we now keep an array to a per cpu structure that--among other things--contains the pointer to the lockless freelist. The freeing thread can then keep possession of exclusive access to the page struct cacheline while the allocating thread keeps its exclusive access to the cacheline containing the per cpu structure. This works as long as the allocating cpu is able to service its request from the lockless freelist. If the lockless freelist runs empty then the allocating thread needs to acquire exclusive access to the cacheline with the page struct lock the slab. The allocating thread will then check if new objects were freed to the per cpu slab. If so it will keep the slab as the cpu slab and continue with the recently remote freed objects. So the allocating thread can take a series of just freed remote pages and dish them out again. Ideally allocations could be just recycling objects in the same slab this way which will lead to an ideal allocation / remote free pattern. The number of objects that can be handled in this way is limited by the capacity of one slab. Increasing slab size via slub_min_objects/ slub_max_order may increase the number of objects and therefore performance. If the allocating thread runs out of objects and finds that no objects were put back by the remote processor then it will retrieve a new slab (from the partial lists or from the page allocator) and start with a whole new set of objects while the remote thread may still be freeing objects to the old cpu slab. This may then repeat until the new slab is also exhausted. If remote freeing has freed objects in the earlier slab then that earlier slab will now be on the partial freelist and the allocating thread will pick that slab next for allocation. So the loop is extended. However, both threads need to take the list_lock to make the swizzling via the partial list happen. It is likely that this kind of scheme will keep the objects being passed around to a small set that can be kept in the cpu caches leading to increased performance. More code cleanups become possible: - Instead of passing a cpu we can now pass a kmem_cache_cpu structure around. Allows reducing the number of parameters to various functions. - Can define a new node_match() function for NUMA to encapsulate locality checks. Effect on allocations: Cachelines touched before this patch: Write: page cache struct and first cacheline of object Cachelines touched after this patch: Write: kmem_cache_cpu cacheline and first cacheline of object Read: page cache struct (but see later patch that avoids touching that cacheline) The handling when the lockless alloc list runs empty gets to be a bit more complicated since another cacheline has now to be written to. But that is halfway out of the hot path. Effect on freeing: Cachelines touched before this patch: Write: page_struct and first cacheline of object Cachelines touched after this patch depending on how we free: Write(to cpu_slab): kmem_cache_cpu struct and first cacheline of object Write(to other): page struct and first cacheline of object Read(to cpu_slab): page struct to id slab etc. (but see later patch that avoids touching the page struct on free) Read(to other): cpu local kmem_cache_cpu struct to verify its not the cpu slab. Summary: Pro: - Distinct cachelines so that concurrent remote frees and local allocs on a cpuslab can occur without cacheline bouncing. - Avoids potential bouncing cachelines because of neighboring per cpu pointer updates in kmem_cache's cpu_slab structure since it now grows to a cacheline (Therefore remove the comment that talks about that concern). Cons: - Freeing objects now requires the reading of one additional cacheline. That can be mitigated for some cases by the following patches but its not possible to completely eliminate these references. - Memory usage grows slightly. The size of each per cpu object is blown up from one word (pointing to the page_struct) to one cacheline with various data. So this is NR_CPUS*NR_SLABS*L1_BYTES more memory use. Lets say NR_SLABS is 100 and a cache line size of 128 then we have just increased SLAB metadata requirements by 12.8k per cpu. (Another later patch reduces these requirements) Signed-off-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 12:26:05 +04:00
struct kmem_cache_cpu {
Lockless (and preemptless) fastpaths for slub Use the this_cpu_cmpxchg_double functionality to implement a lockless allocation algorithm on arches that support fast this_cpu_ops. Each of the per cpu pointers is paired with a transaction id that ensures that updates of the per cpu information can only occur in sequence on a certain cpu. A transaction id is a "long" integer that is comprised of an event number and the cpu number. The event number is incremented for every change to the per cpu state. This means that the cmpxchg instruction can verify for an update that nothing interfered and that we are updating the percpu structure for the processor where we picked up the information and that we are also currently on that processor when we update the information. This results in a significant decrease of the overhead in the fastpaths. It also makes it easy to adopt the fast path for realtime kernels since this is lockless and does not require the use of the current per cpu area over the critical section. It is only important that the per cpu area is current at the beginning of the critical section and at the end. So there is no need even to disable preemption. Test results show that the fastpath cycle count is reduced by up to ~ 40% (alloc/free test goes from ~140 cycles down to ~80). The slowpath for kfree adds a few cycles. Sadly this does nothing for the slowpath which is where the main issues with performance in slub are but the best case performance rises significantly. (For that see the more complex slub patches that require cmpxchg_double) Kmalloc: alloc/free test Before: 10000 times kmalloc(8)/kfree -> 134 cycles 10000 times kmalloc(16)/kfree -> 152 cycles 10000 times kmalloc(32)/kfree -> 144 cycles 10000 times kmalloc(64)/kfree -> 142 cycles 10000 times kmalloc(128)/kfree -> 142 cycles 10000 times kmalloc(256)/kfree -> 132 cycles 10000 times kmalloc(512)/kfree -> 132 cycles 10000 times kmalloc(1024)/kfree -> 135 cycles 10000 times kmalloc(2048)/kfree -> 135 cycles 10000 times kmalloc(4096)/kfree -> 135 cycles 10000 times kmalloc(8192)/kfree -> 144 cycles 10000 times kmalloc(16384)/kfree -> 754 cycles After: 10000 times kmalloc(8)/kfree -> 78 cycles 10000 times kmalloc(16)/kfree -> 78 cycles 10000 times kmalloc(32)/kfree -> 82 cycles 10000 times kmalloc(64)/kfree -> 88 cycles 10000 times kmalloc(128)/kfree -> 79 cycles 10000 times kmalloc(256)/kfree -> 79 cycles 10000 times kmalloc(512)/kfree -> 85 cycles 10000 times kmalloc(1024)/kfree -> 82 cycles 10000 times kmalloc(2048)/kfree -> 82 cycles 10000 times kmalloc(4096)/kfree -> 85 cycles 10000 times kmalloc(8192)/kfree -> 82 cycles 10000 times kmalloc(16384)/kfree -> 706 cycles Kmalloc: Repeatedly allocate then free test Before: 10000 times kmalloc(8) -> 211 cycles kfree -> 113 cycles 10000 times kmalloc(16) -> 174 cycles kfree -> 115 cycles 10000 times kmalloc(32) -> 235 cycles kfree -> 129 cycles 10000 times kmalloc(64) -> 222 cycles kfree -> 120 cycles 10000 times kmalloc(128) -> 343 cycles kfree -> 139 cycles 10000 times kmalloc(256) -> 827 cycles kfree -> 147 cycles 10000 times kmalloc(512) -> 1048 cycles kfree -> 272 cycles 10000 times kmalloc(1024) -> 2043 cycles kfree -> 528 cycles 10000 times kmalloc(2048) -> 4002 cycles kfree -> 571 cycles 10000 times kmalloc(4096) -> 7740 cycles kfree -> 628 cycles 10000 times kmalloc(8192) -> 8062 cycles kfree -> 850 cycles 10000 times kmalloc(16384) -> 8895 cycles kfree -> 1249 cycles After: 10000 times kmalloc(8) -> 190 cycles kfree -> 129 cycles 10000 times kmalloc(16) -> 76 cycles kfree -> 123 cycles 10000 times kmalloc(32) -> 126 cycles kfree -> 124 cycles 10000 times kmalloc(64) -> 181 cycles kfree -> 128 cycles 10000 times kmalloc(128) -> 310 cycles kfree -> 140 cycles 10000 times kmalloc(256) -> 809 cycles kfree -> 165 cycles 10000 times kmalloc(512) -> 1005 cycles kfree -> 269 cycles 10000 times kmalloc(1024) -> 1999 cycles kfree -> 527 cycles 10000 times kmalloc(2048) -> 3967 cycles kfree -> 570 cycles 10000 times kmalloc(4096) -> 7658 cycles kfree -> 637 cycles 10000 times kmalloc(8192) -> 8111 cycles kfree -> 859 cycles 10000 times kmalloc(16384) -> 8791 cycles kfree -> 1173 cycles Signed-off-by: Christoph Lameter <cl@linux.com> Signed-off-by: Pekka Enberg <penberg@kernel.org>
2011-02-25 20:38:54 +03:00
void **freelist; /* Pointer to next available object */
unsigned long tid; /* Globally unique transaction id */
struct page *page; /* The slab from which we are allocating */
#ifdef CONFIG_SLUB_CPU_PARTIAL
slub: per cpu cache for partial pages Allow filling out the rest of the kmem_cache_cpu cacheline with pointers to partial pages. The partial page list is used in slab_free() to avoid per node lock taking. In __slab_alloc() we can then take multiple partial pages off the per node partial list in one go reducing node lock pressure. We can also use the per cpu partial list in slab_alloc() to avoid scanning partial lists for pages with free objects. The main effect of a per cpu partial list is that the per node list_lock is taken for batches of partial pages instead of individual ones. Potential future enhancements: 1. The pickup from the partial list could be perhaps be done without disabling interrupts with some work. The free path already puts the page into the per cpu partial list without disabling interrupts. 2. __slab_free() may have some code paths that could use optimization. Performance: Before After ./hackbench 100 process 200000 Time: 1953.047 1564.614 ./hackbench 100 process 20000 Time: 207.176 156.940 ./hackbench 100 process 20000 Time: 204.468 156.940 ./hackbench 100 process 20000 Time: 204.879 158.772 ./hackbench 10 process 20000 Time: 20.153 15.853 ./hackbench 10 process 20000 Time: 20.153 15.986 ./hackbench 10 process 20000 Time: 19.363 16.111 ./hackbench 1 process 20000 Time: 2.518 2.307 ./hackbench 1 process 20000 Time: 2.258 2.339 ./hackbench 1 process 20000 Time: 2.864 2.163 Signed-off-by: Christoph Lameter <cl@linux.com> Signed-off-by: Pekka Enberg <penberg@kernel.org>
2011-08-10 01:12:27 +04:00
struct page *partial; /* Partially allocated frozen slabs */
#endif
SLUB: Support for performance statistics The statistics provided here allow the monitoring of allocator behavior but at the cost of some (minimal) loss of performance. Counters are placed in SLUB's per cpu data structure. The per cpu structure may be extended by the statistics to grow larger than one cacheline which will increase the cache footprint of SLUB. There is a compile option to enable/disable the inclusion of the runtime statistics and its off by default. The slabinfo tool is enhanced to support these statistics via two options: -D Switches the line of information displayed for a slab from size mode to activity mode. -A Sorts the slabs displayed by activity. This allows the display of the slabs most important to the performance of a certain load. -r Report option will report detailed statistics on Example (tbench load): slabinfo -AD ->Shows the most active slabs Name Objects Alloc Free %Fast skbuff_fclone_cache 33 111953835 111953835 99 99 :0000192 2666 5283688 5281047 99 99 :0001024 849 5247230 5246389 83 83 vm_area_struct 1349 119642 118355 91 22 :0004096 15 66753 66751 98 98 :0000064 2067 25297 23383 98 78 dentry 10259 28635 18464 91 45 :0000080 11004 18950 8089 98 98 :0000096 1703 12358 10784 99 98 :0000128 762 10582 9875 94 18 :0000512 184 9807 9647 95 81 :0002048 479 9669 9195 83 65 anon_vma 777 9461 9002 99 71 kmalloc-8 6492 9981 5624 99 97 :0000768 258 7174 6931 58 15 So the skbuff_fclone_cache is of highest importance for the tbench load. Pretty high load on the 192 sized slab. Look for the aliases slabinfo -a | grep 000192 :0000192 <- xfs_btree_cur filp kmalloc-192 uid_cache tw_sock_TCP request_sock_TCPv6 tw_sock_TCPv6 skbuff_head_cache xfs_ili Likely skbuff_head_cache. Looking into the statistics of the skbuff_fclone_cache is possible through slabinfo skbuff_fclone_cache ->-r option implied if cache name is mentioned .... Usual output ... Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 111953360 111946981 99 99 Slowpath 1044 7423 0 0 Page Alloc 272 264 0 0 Add partial 25 325 0 0 Remove partial 86 264 0 0 RemoteObj/SlabFrozen 350 4832 0 0 Total 111954404 111954404 Flushes 49 Refill 0 Deactivate Full=325(92%) Empty=0(0%) ToHead=24(6%) ToTail=1(0%) Looks good because the fastpath is overwhelmingly taken. skbuff_head_cache: Slab Perf Counter Alloc Free %Al %Fr -------------------------------------------------- Fastpath 5297262 5259882 99 99 Slowpath 4477 39586 0 0 Page Alloc 937 824 0 0 Add partial 0 2515 0 0 Remove partial 1691 824 0 0 RemoteObj/SlabFrozen 2621 9684 0 0 Total 5301739 5299468 Deactivate Full=2620(100%) Empty=0(0%) ToHead=0(0%) ToTail=0(0%) Descriptions of the output: Total: The total number of allocation and frees that occurred for a slab Fastpath: The number of allocations/frees that used the fastpath. Slowpath: Other allocations Page Alloc: Number of calls to the page allocator as a result of slowpath processing Add Partial: Number of slabs added to the partial list through free or alloc (occurs during cpuslab flushes) Remove Partial: Number of slabs removed from the partial list as a result of allocations retrieving a partial slab or by a free freeing the last object of a slab. RemoteObj/Froz: How many times were remotely freed object encountered when a slab was about to be deactivated. Frozen: How many times was free able to skip list processing because the slab was in use as the cpuslab of another processor. Flushes: Number of times the cpuslab was flushed on request (kmem_cache_shrink, may result from races in __slab_alloc) Refill: Number of times we were able to refill the cpuslab from remotely freed objects for the same slab. Deactivate: Statistics how slabs were deactivated. Shows how they were put onto the partial list. In general fastpath is very good. Slowpath without partial list processing is also desirable. Any touching of partial list uses node specific locks which may potentially cause list lock contention. Signed-off-by: Christoph Lameter <clameter@sgi.com>
2008-02-08 04:47:41 +03:00
#ifdef CONFIG_SLUB_STATS
unsigned stat[NR_SLUB_STAT_ITEMS];
#endif
SLUB: Place kmem_cache_cpu structures in a NUMA aware way The kmem_cache_cpu structures introduced are currently an array placed in the kmem_cache struct. Meaning the kmem_cache_cpu structures are overwhelmingly on the wrong node for systems with a higher amount of nodes. These are performance critical structures since the per node information has to be touched for every alloc and free in a slab. In order to place the kmem_cache_cpu structure optimally we put an array of pointers to kmem_cache_cpu structs in kmem_cache (similar to SLAB). However, the kmem_cache_cpu structures can now be allocated in a more intelligent way. We would like to put per cpu structures for the same cpu but different slab caches in cachelines together to save space and decrease the cache footprint. However, the slab allocators itself control only allocations per node. We set up a simple per cpu array for every processor with 100 per cpu structures which is usually enough to get them all set up right. If we run out then we fall back to kmalloc_node. This also solves the bootstrap problem since we do not have to use slab allocator functions early in boot to get memory for the small per cpu structures. Pro: - NUMA aware placement improves memory performance - All global structures in struct kmem_cache become readonly - Dense packing of per cpu structures reduces cacheline footprint in SMP and NUMA. - Potential avoidance of exclusive cacheline fetches on the free and alloc hotpath since multiple kmem_cache_cpu structures are in one cacheline. This is particularly important for the kmalloc array. Cons: - Additional reference to one read only cacheline (per cpu array of pointers to kmem_cache_cpu) in both slab_alloc() and slab_free(). [akinobu.mita@gmail.com: fix cpu hotplug offline/online path] Signed-off-by: Christoph Lameter <clameter@sgi.com> Cc: "Pekka Enberg" <penberg@cs.helsinki.fi> Cc: Akinobu Mita <akinobu.mita@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-16 12:26:08 +04:00
};
SLUB: Avoid page struct cacheline bouncing due to remote frees to cpu slab A remote free may access the same page struct that also contains the lockless freelist for the cpu slab. If objects have a short lifetime and are freed by a different processor then remote frees back to the slab from which we are currently allocating are frequent. The cacheline with the page struct needs to be repeately acquired in exclusive mode by both the allocating thread and the freeing thread. If this is frequent enough then performance will suffer because of cacheline bouncing. This patchset puts the lockless_freelist pointer in its own cacheline. In order to make that happen we introduce a per cpu structure called kmem_cache_cpu. Instead of keeping an array of pointers to page structs we now keep an array to a per cpu structure that--among other things--contains the pointer to the lockless freelist. The freeing thread can then keep possession of exclusive access to the page struct cacheline while the allocating thread keeps its exclusive access to the cacheline containing the per cpu structure. This works as long as the allocating cpu is able to service its request from the lockless freelist. If the lockless freelist runs empty then the allocating thread needs to acquire exclusive access to the cacheline with the page struct lock the slab. The allocating thread will then check if new objects were freed to the per cpu slab. If so it will keep the slab as the cpu slab and continue with the recently remote freed objects. So the allocating thread can take a series of just freed remote pages and dish them out again. Ideally allocations could be just recycling objects in the same slab this way which will lead to an ideal allocation / remote free pattern. The number of objects that can be handled in this way is limited by the capacity of one slab. Increasing slab size via slub_min_objects/ slub_max_order may increase the number of objects and therefore performance. If the allocating thread runs out of objects and finds that no objects were put back by the remote processor then it will retrieve a new slab (from the partial lists or from the page allocator) and start with a whole new set of objects while the remote thread may still be freeing objects to the old cpu slab. This may then repeat until the new slab is also exhausted. If remote freeing has freed objects in the earlier slab then that earlier slab will now be on the partial freelist and the allocating thread will pick that slab next for allocation. So the loop is extended. However, both threads need to take the list_lock to make the swizzling via the partial list happen. It is likely that this kind of scheme will keep the objects being passed around to a small set that can be kept in the cpu caches leading to increased performance. More code cleanups become possible: - Instead of passing a cpu we can now pass a kmem_cache_cpu structure around. Allows reducing the number of parameters to various functions. - Can define a new node_match() function for NUMA to encapsulate locality checks. Effect on allocations: Cachelines touched before this patch: Write: page cache struct and first cacheline of object Cachelines touched after this patch: Write: kmem_cache_cpu cacheline and first cacheline of object Read: page cache struct (but see later patch that avoids touching that cacheline) The handling when the lockless alloc list runs empty gets to be a bit more complicated since another cacheline has now to be written to. But that is halfway out of the hot path. Effect on freeing: Cachelines touched before this patch: Write: page_struct and first cacheline of object Cachelines touched after this patch depending on how we free: Write(to cpu_slab): kmem_cache_cpu struct and first cacheline of object Write(to other): page struct and first cacheline of object Read(to cpu_slab): page struct to id slab etc. (but see later patch that avoids touching the page struct on free) Read(to other): cpu local kmem_cache_cpu struct to verify its not the cpu slab. Summary: Pro: - Distinct cachelines so that concurrent remote frees and local allocs on a cpuslab can occur without cacheline bouncing. - Avoids potential bouncing cachelines because of neighboring per cpu pointer updates in kmem_cache's cpu_slab structure since it now grows to a cacheline (Therefore remove the comment that talks about that concern). Cons: - Freeing objects now requires the reading of one additional cacheline. That can be mitigated for some cases by the following patches but its not possible to completely eliminate these references. - Memory usage grows slightly. The size of each per cpu object is blown up from one word (pointing to the page_struct) to one cacheline with various data. So this is NR_CPUS*NR_SLABS*L1_BYTES more memory use. Lets say NR_SLABS is 100 and a cache line size of 128 then we have just increased SLAB metadata requirements by 12.8k per cpu. (Another later patch reduces these requirements) Signed-off-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 12:26:05 +04:00
#ifdef CONFIG_SLUB_CPU_PARTIAL
#define slub_percpu_partial(c) ((c)->partial)
#define slub_set_percpu_partial(c, p) \
({ \
slub_percpu_partial(c) = (p)->next; \
})
#define slub_percpu_partial_read_once(c) READ_ONCE(slub_percpu_partial(c))
#else
#define slub_percpu_partial(c) NULL
#define slub_set_percpu_partial(c, p)
#define slub_percpu_partial_read_once(c) NULL
#endif // CONFIG_SLUB_CPU_PARTIAL
/*
* Word size structure that can be atomically updated or read and that
* contains both the order and the number of objects that a slab of the
* given order would contain.
*/
struct kmem_cache_order_objects {
unsigned int x;
};
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
/*
* Slab cache management.
*/
struct kmem_cache {
struct kmem_cache_cpu __percpu *cpu_slab;
/* Used for retrieving partial slabs, etc. */
slab_flags_t flags;
unsigned long min_partial;
unsigned int size; /* The size of an object including metadata */
unsigned int object_size;/* The size of an object without metadata */
unsigned int offset; /* Free pointer offset */
#ifdef CONFIG_SLUB_CPU_PARTIAL
/* Number of per cpu partial objects to keep around */
unsigned int cpu_partial;
#endif
struct kmem_cache_order_objects oo;
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
/* Allocation and freeing of slabs */
struct kmem_cache_order_objects max;
struct kmem_cache_order_objects min;
gfp_t allocflags; /* gfp flags to use on each alloc */
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
int refcount; /* Refcount for slab cache destroy */
void (*ctor)(void *);
unsigned int inuse; /* Offset to metadata */
unsigned int align; /* Alignment */
unsigned int red_left_pad; /* Left redzone padding size */
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
const char *name; /* Name (only for display!) */
struct list_head list; /* List of slab caches */
#ifdef CONFIG_SYSFS
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
struct kobject kobj; /* For sysfs */
slub: make sysfs file removal asynchronous Commit bf5eb3de3847 ("slub: separate out sysfs_slab_release() from sysfs_slab_remove()") made slub sysfs file removals synchronous to kmem_cache shutdown. Unfortunately, this created a possible ABBA deadlock between slab_mutex and sysfs draining mechanism triggering the following lockdep warning. ====================================================== [ INFO: possible circular locking dependency detected ] 4.10.0-test+ #48 Not tainted ------------------------------------------------------- rmmod/1211 is trying to acquire lock: (s_active#120){++++.+}, at: [<ffffffff81308073>] kernfs_remove+0x23/0x40 but task is already holding lock: (slab_mutex){+.+.+.}, at: [<ffffffff8120f691>] kmem_cache_destroy+0x41/0x2d0 which lock already depends on the new lock. the existing dependency chain (in reverse order) is: -> #1 (slab_mutex){+.+.+.}: lock_acquire+0xf6/0x1f0 __mutex_lock+0x75/0x950 mutex_lock_nested+0x1b/0x20 slab_attr_store+0x75/0xd0 sysfs_kf_write+0x45/0x60 kernfs_fop_write+0x13c/0x1c0 __vfs_write+0x28/0x120 vfs_write+0xc8/0x1e0 SyS_write+0x49/0xa0 entry_SYSCALL_64_fastpath+0x1f/0xc2 -> #0 (s_active#120){++++.+}: __lock_acquire+0x10ed/0x1260 lock_acquire+0xf6/0x1f0 __kernfs_remove+0x254/0x320 kernfs_remove+0x23/0x40 sysfs_remove_dir+0x51/0x80 kobject_del+0x18/0x50 __kmem_cache_shutdown+0x3e6/0x460 kmem_cache_destroy+0x1fb/0x2d0 kvm_exit+0x2d/0x80 [kvm] vmx_exit+0x19/0xa1b [kvm_intel] SyS_delete_module+0x198/0x1f0 entry_SYSCALL_64_fastpath+0x1f/0xc2 other info that might help us debug this: Possible unsafe locking scenario: CPU0 CPU1 ---- ---- lock(slab_mutex); lock(s_active#120); lock(slab_mutex); lock(s_active#120); *** DEADLOCK *** 2 locks held by rmmod/1211: #0: (cpu_hotplug.dep_map){++++++}, at: [<ffffffff810a7877>] get_online_cpus+0x37/0x80 #1: (slab_mutex){+.+.+.}, at: [<ffffffff8120f691>] kmem_cache_destroy+0x41/0x2d0 stack backtrace: CPU: 3 PID: 1211 Comm: rmmod Not tainted 4.10.0-test+ #48 Hardware name: Hewlett-Packard HP Compaq Pro 6300 SFF/339A, BIOS K01 v02.05 05/07/2012 Call Trace: print_circular_bug+0x1be/0x210 __lock_acquire+0x10ed/0x1260 lock_acquire+0xf6/0x1f0 __kernfs_remove+0x254/0x320 kernfs_remove+0x23/0x40 sysfs_remove_dir+0x51/0x80 kobject_del+0x18/0x50 __kmem_cache_shutdown+0x3e6/0x460 kmem_cache_destroy+0x1fb/0x2d0 kvm_exit+0x2d/0x80 [kvm] vmx_exit+0x19/0xa1b [kvm_intel] SyS_delete_module+0x198/0x1f0 ? SyS_delete_module+0x5/0x1f0 entry_SYSCALL_64_fastpath+0x1f/0xc2 It'd be the cleanest to deal with the issue by removing sysfs files without holding slab_mutex before the rest of shutdown; however, given the current code structure, it is pretty difficult to do so. This patch punts sysfs file removal to a work item. Before commit bf5eb3de3847, the removal was punted to a RCU delayed work item which is executed after release. Now, we're punting to a different work item on shutdown which still maintains the goal removing the sysfs files earlier when destroying kmem_caches. Link: http://lkml.kernel.org/r/20170620204512.GI21326@htj.duckdns.org Fixes: bf5eb3de3847 ("slub: separate out sysfs_slab_release() from sysfs_slab_remove()") Signed-off-by: Tejun Heo <tj@kernel.org> Reported-by: Steven Rostedt (VMware) <rostedt@goodmis.org> Tested-by: Steven Rostedt (VMware) <rostedt@goodmis.org> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-06-24 01:08:52 +03:00
struct work_struct kobj_remove_work;
#endif
#ifdef CONFIG_MEMCG
slab: embed memcg_cache_params to kmem_cache Currently, kmem_cache stores a pointer to struct memcg_cache_params instead of embedding it. The rationale is to save memory when kmem accounting is disabled. However, the memcg_cache_params has shrivelled drastically since it was first introduced: * Initially: struct memcg_cache_params { bool is_root_cache; union { struct kmem_cache *memcg_caches[0]; struct { struct mem_cgroup *memcg; struct list_head list; struct kmem_cache *root_cache; bool dead; atomic_t nr_pages; struct work_struct destroy; }; }; }; * Now: struct memcg_cache_params { bool is_root_cache; union { struct { struct rcu_head rcu_head; struct kmem_cache *memcg_caches[0]; }; struct { struct mem_cgroup *memcg; struct kmem_cache *root_cache; }; }; }; So the memory saving does not seem to be a clear win anymore. OTOH, keeping a pointer to memcg_cache_params struct instead of embedding it results in touching one more cache line on kmem alloc/free hot paths. Besides, it makes linking kmem caches in a list chained by a field of struct memcg_cache_params really painful due to a level of indirection, while I want to make them linked in the following patch. That said, let us embed it. Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Dave Chinner <david@fromorbit.com> Cc: Dan Carpenter <dan.carpenter@oracle.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-13 01:59:20 +03:00
struct memcg_cache_params memcg_params;
/* For propagation, maximum size of a stored attr */
unsigned int max_attr_size;
2014-04-08 02:39:31 +04:00
#ifdef CONFIG_SYSFS
struct kset *memcg_kset;
#endif
#endif
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
mm: add SLUB free list pointer obfuscation This SLUB free list pointer obfuscation code is modified from Brad Spengler/PaX Team's code in the last public patch of grsecurity/PaX based on my understanding of the code. Changes or omissions from the original code are mine and don't reflect the original grsecurity/PaX code. This adds a per-cache random value to SLUB caches that is XORed with their freelist pointer address and value. This adds nearly zero overhead and frustrates the very common heap overflow exploitation method of overwriting freelist pointers. A recent example of the attack is written up here: http://cyseclabs.com/blog/cve-2016-6187-heap-off-by-one-exploit and there is a section dedicated to the technique the book "A Guide to Kernel Exploitation: Attacking the Core". This is based on patches by Daniel Micay, and refactored to minimize the use of #ifdef. With 200-count cycles of "hackbench -g 20 -l 1000" I saw the following run times: before: mean 10.11882499999999999995 variance .03320378329145728642 stdev .18221905304181911048 after: mean 10.12654000000000000014 variance .04700556623115577889 stdev .21680767106160192064 The difference gets lost in the noise, but if the above is to be taken literally, using CONFIG_FREELIST_HARDENED is 0.07% slower. Link: http://lkml.kernel.org/r/20170802180609.GA66807@beast Signed-off-by: Kees Cook <keescook@chromium.org> Suggested-by: Daniel Micay <danielmicay@gmail.com> Cc: Rik van Riel <riel@redhat.com> Cc: Tycho Andersen <tycho@docker.com> Cc: Alexander Popov <alex.popov@linux.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-09-07 02:19:18 +03:00
#ifdef CONFIG_SLAB_FREELIST_HARDENED
unsigned long random;
#endif
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
#ifdef CONFIG_NUMA
/*
* Defragmentation by allocating from a remote node.
*/
unsigned int remote_node_defrag_ratio;
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
#endif
mm: SLUB freelist randomization Implements freelist randomization for the SLUB allocator. It was previous implemented for the SLAB allocator. Both use the same configuration option (CONFIG_SLAB_FREELIST_RANDOM). The list is randomized during initialization of a new set of pages. The order on different freelist sizes is pre-computed at boot for performance. Each kmem_cache has its own randomized freelist. This security feature reduces the predictability of the kernel SLUB allocator against heap overflows rendering attacks much less stable. For example these attacks exploit the predictability of the heap: - Linux Kernel CAN SLUB overflow (https://goo.gl/oMNWkU) - Exploiting Linux Kernel Heap corruptions (http://goo.gl/EXLn95) Performance results: slab_test impact is between 3% to 4% on average for 100000 attempts without smp. It is a very focused testing, kernbench show the overall impact on the system is way lower. Before: Single thread testing ===================== 1. Kmalloc: Repeatedly allocate then free test 100000 times kmalloc(8) -> 49 cycles kfree -> 77 cycles 100000 times kmalloc(16) -> 51 cycles kfree -> 79 cycles 100000 times kmalloc(32) -> 53 cycles kfree -> 83 cycles 100000 times kmalloc(64) -> 62 cycles kfree -> 90 cycles 100000 times kmalloc(128) -> 81 cycles kfree -> 97 cycles 100000 times kmalloc(256) -> 98 cycles kfree -> 121 cycles 100000 times kmalloc(512) -> 95 cycles kfree -> 122 cycles 100000 times kmalloc(1024) -> 96 cycles kfree -> 126 cycles 100000 times kmalloc(2048) -> 115 cycles kfree -> 140 cycles 100000 times kmalloc(4096) -> 149 cycles kfree -> 171 cycles 2. Kmalloc: alloc/free test 100000 times kmalloc(8)/kfree -> 70 cycles 100000 times kmalloc(16)/kfree -> 70 cycles 100000 times kmalloc(32)/kfree -> 70 cycles 100000 times kmalloc(64)/kfree -> 70 cycles 100000 times kmalloc(128)/kfree -> 70 cycles 100000 times kmalloc(256)/kfree -> 69 cycles 100000 times kmalloc(512)/kfree -> 70 cycles 100000 times kmalloc(1024)/kfree -> 73 cycles 100000 times kmalloc(2048)/kfree -> 72 cycles 100000 times kmalloc(4096)/kfree -> 71 cycles After: Single thread testing ===================== 1. Kmalloc: Repeatedly allocate then free test 100000 times kmalloc(8) -> 57 cycles kfree -> 78 cycles 100000 times kmalloc(16) -> 61 cycles kfree -> 81 cycles 100000 times kmalloc(32) -> 76 cycles kfree -> 93 cycles 100000 times kmalloc(64) -> 83 cycles kfree -> 94 cycles 100000 times kmalloc(128) -> 106 cycles kfree -> 107 cycles 100000 times kmalloc(256) -> 118 cycles kfree -> 117 cycles 100000 times kmalloc(512) -> 114 cycles kfree -> 116 cycles 100000 times kmalloc(1024) -> 115 cycles kfree -> 118 cycles 100000 times kmalloc(2048) -> 147 cycles kfree -> 131 cycles 100000 times kmalloc(4096) -> 214 cycles kfree -> 161 cycles 2. Kmalloc: alloc/free test 100000 times kmalloc(8)/kfree -> 66 cycles 100000 times kmalloc(16)/kfree -> 66 cycles 100000 times kmalloc(32)/kfree -> 66 cycles 100000 times kmalloc(64)/kfree -> 66 cycles 100000 times kmalloc(128)/kfree -> 65 cycles 100000 times kmalloc(256)/kfree -> 67 cycles 100000 times kmalloc(512)/kfree -> 67 cycles 100000 times kmalloc(1024)/kfree -> 64 cycles 100000 times kmalloc(2048)/kfree -> 67 cycles 100000 times kmalloc(4096)/kfree -> 67 cycles Kernbench, before: Average Optimal load -j 12 Run (std deviation): Elapsed Time 101.873 (1.16069) User Time 1045.22 (1.60447) System Time 88.969 (0.559195) Percent CPU 1112.9 (13.8279) Context Switches 189140 (2282.15) Sleeps 99008.6 (768.091) After: Average Optimal load -j 12 Run (std deviation): Elapsed Time 102.47 (0.562732) User Time 1045.3 (1.34263) System Time 88.311 (0.342554) Percent CPU 1105.8 (6.49444) Context Switches 189081 (2355.78) Sleeps 99231.5 (800.358) Link: http://lkml.kernel.org/r/1464295031-26375-3-git-send-email-thgarnie@google.com Signed-off-by: Thomas Garnier <thgarnie@google.com> Reviewed-by: Kees Cook <keescook@chromium.org> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-27 01:21:59 +03:00
#ifdef CONFIG_SLAB_FREELIST_RANDOM
unsigned int *random_seq;
#endif
#ifdef CONFIG_KASAN
struct kasan_cache kasan_info;
#endif
unsigned int useroffset; /* Usercopy region offset */
unsigned int usersize; /* Usercopy region size */
usercopy: Prepare for usercopy whitelisting This patch prepares the slab allocator to handle caches having annotations (useroffset and usersize) defining usercopy regions. This patch is modified from Brad Spengler/PaX Team's PAX_USERCOPY whitelisting code in the last public patch of grsecurity/PaX based on my understanding of the code. Changes or omissions from the original code are mine and don't reflect the original grsecurity/PaX code. Currently, hardened usercopy performs dynamic bounds checking on slab cache objects. This is good, but still leaves a lot of kernel memory available to be copied to/from userspace in the face of bugs. To further restrict what memory is available for copying, this creates a way to whitelist specific areas of a given slab cache object for copying to/from userspace, allowing much finer granularity of access control. Slab caches that are never exposed to userspace can declare no whitelist for their objects, thereby keeping them unavailable to userspace via dynamic copy operations. (Note, an implicit form of whitelisting is the use of constant sizes in usercopy operations and get_user()/put_user(); these bypass hardened usercopy checks since these sizes cannot change at runtime.) To support this whitelist annotation, usercopy region offset and size members are added to struct kmem_cache. The slab allocator receives a new function, kmem_cache_create_usercopy(), that creates a new cache with a usercopy region defined, suitable for declaring spans of fields within the objects that get copied to/from userspace. In this patch, the default kmem_cache_create() marks the entire allocation as whitelisted, leaving it semantically unchanged. Once all fine-grained whitelists have been added (in subsequent patches), this will be changed to a usersize of 0, making caches created with kmem_cache_create() not copyable to/from userspace. After the entire usercopy whitelist series is applied, less than 15% of the slab cache memory remains exposed to potential usercopy bugs after a fresh boot: Total Slab Memory: 48074720 Usercopyable Memory: 6367532 13.2% task_struct 0.2% 4480/1630720 RAW 0.3% 300/96000 RAWv6 2.1% 1408/64768 ext4_inode_cache 3.0% 269760/8740224 dentry 11.1% 585984/5273856 mm_struct 29.1% 54912/188448 kmalloc-8 100.0% 24576/24576 kmalloc-16 100.0% 28672/28672 kmalloc-32 100.0% 81920/81920 kmalloc-192 100.0% 96768/96768 kmalloc-128 100.0% 143360/143360 names_cache 100.0% 163840/163840 kmalloc-64 100.0% 167936/167936 kmalloc-256 100.0% 339968/339968 kmalloc-512 100.0% 350720/350720 kmalloc-96 100.0% 455616/455616 kmalloc-8192 100.0% 655360/655360 kmalloc-1024 100.0% 812032/812032 kmalloc-4096 100.0% 819200/819200 kmalloc-2048 100.0% 1310720/1310720 After some kernel build workloads, the percentage (mainly driven by dentry and inode caches expanding) drops under 10%: Total Slab Memory: 95516184 Usercopyable Memory: 8497452 8.8% task_struct 0.2% 4000/1456000 RAW 0.3% 300/96000 RAWv6 2.1% 1408/64768 ext4_inode_cache 3.0% 1217280/39439872 dentry 11.1% 1623200/14608800 mm_struct 29.1% 73216/251264 kmalloc-8 100.0% 24576/24576 kmalloc-16 100.0% 28672/28672 kmalloc-32 100.0% 94208/94208 kmalloc-192 100.0% 96768/96768 kmalloc-128 100.0% 143360/143360 names_cache 100.0% 163840/163840 kmalloc-64 100.0% 245760/245760 kmalloc-256 100.0% 339968/339968 kmalloc-512 100.0% 350720/350720 kmalloc-96 100.0% 563520/563520 kmalloc-8192 100.0% 655360/655360 kmalloc-1024 100.0% 794624/794624 kmalloc-4096 100.0% 819200/819200 kmalloc-2048 100.0% 1257472/1257472 Signed-off-by: David Windsor <dave@nullcore.net> [kees: adjust commit log, split out a few extra kmalloc hunks] [kees: add field names to function declarations] [kees: convert BUGs to WARNs and fail closed] [kees: add attack surface reduction analysis to commit log] Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: linux-mm@kvack.org Cc: linux-xfs@vger.kernel.org Signed-off-by: Kees Cook <keescook@chromium.org> Acked-by: Christoph Lameter <cl@linux.com>
2017-06-11 05:50:28 +03:00
struct kmem_cache_node *node[MAX_NUMNODES];
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
};
#ifdef CONFIG_SLUB_CPU_PARTIAL
#define slub_cpu_partial(s) ((s)->cpu_partial)
#define slub_set_cpu_partial(s, n) \
({ \
slub_cpu_partial(s) = (n); \
})
#else
#define slub_cpu_partial(s) (0)
#define slub_set_cpu_partial(s, n)
#endif /* CONFIG_SLUB_CPU_PARTIAL */
slub: use sysfs'es release mechanism for kmem_cache debugobjects warning during netfilter exit: ------------[ cut here ]------------ WARNING: CPU: 6 PID: 4178 at lib/debugobjects.c:260 debug_print_object+0x8d/0xb0() ODEBUG: free active (active state 0) object type: timer_list hint: delayed_work_timer_fn+0x0/0x20 Modules linked in: CPU: 6 PID: 4178 Comm: kworker/u16:2 Tainted: G W 3.11.0-next-20130906-sasha #3984 Workqueue: netns cleanup_net Call Trace: dump_stack+0x52/0x87 warn_slowpath_common+0x8c/0xc0 warn_slowpath_fmt+0x46/0x50 debug_print_object+0x8d/0xb0 __debug_check_no_obj_freed+0xa5/0x220 debug_check_no_obj_freed+0x15/0x20 kmem_cache_free+0x197/0x340 kmem_cache_destroy+0x86/0xe0 nf_conntrack_cleanup_net_list+0x131/0x170 nf_conntrack_pernet_exit+0x5d/0x70 ops_exit_list+0x5e/0x70 cleanup_net+0xfb/0x1c0 process_one_work+0x338/0x550 worker_thread+0x215/0x350 kthread+0xe7/0xf0 ret_from_fork+0x7c/0xb0 Also during dcookie cleanup: WARNING: CPU: 12 PID: 9725 at lib/debugobjects.c:260 debug_print_object+0x8c/0xb0() ODEBUG: free active (active state 0) object type: timer_list hint: delayed_work_timer_fn+0x0/0x20 Modules linked in: CPU: 12 PID: 9725 Comm: trinity-c141 Not tainted 3.15.0-rc2-next-20140423-sasha-00018-gc4ff6c4 #408 Call Trace: dump_stack (lib/dump_stack.c:52) warn_slowpath_common (kernel/panic.c:430) warn_slowpath_fmt (kernel/panic.c:445) debug_print_object (lib/debugobjects.c:262) __debug_check_no_obj_freed (lib/debugobjects.c:697) debug_check_no_obj_freed (lib/debugobjects.c:726) kmem_cache_free (mm/slub.c:2689 mm/slub.c:2717) kmem_cache_destroy (mm/slab_common.c:363) dcookie_unregister (fs/dcookies.c:302 fs/dcookies.c:343) event_buffer_release (arch/x86/oprofile/../../../drivers/oprofile/event_buffer.c:153) __fput (fs/file_table.c:217) ____fput (fs/file_table.c:253) task_work_run (kernel/task_work.c:125 (discriminator 1)) do_notify_resume (include/linux/tracehook.h:196 arch/x86/kernel/signal.c:751) int_signal (arch/x86/kernel/entry_64.S:807) Sysfs has a release mechanism. Use that to release the kmem_cache structure if CONFIG_SYSFS is enabled. Only slub is changed - slab currently only supports /proc/slabinfo and not /sys/kernel/slab/*. We talked about adding that and someone was working on it. [akpm@linux-foundation.org: fix CONFIG_SYSFS=n build] [akpm@linux-foundation.org: fix CONFIG_SYSFS=n build even more] Signed-off-by: Christoph Lameter <cl@linux.com> Reported-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Acked-by: Greg KH <greg@kroah.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Pekka Enberg <penberg@kernel.org> Cc: Russell King <rmk@arm.linux.org.uk> Cc: Bart Van Assche <bvanassche@acm.org> Cc: Al Viro <viro@ZenIV.linux.org.uk> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-05-06 23:50:08 +04:00
#ifdef CONFIG_SYSFS
#define SLAB_SUPPORTS_SYSFS
slub: fix failure when we delete and create a slab cache In kernel 4.17 I removed some code from dm-bufio that did slab cache merging (commit 21bb13276768: "dm bufio: remove code that merges slab caches") - both slab and slub support merging caches with identical attributes, so dm-bufio now just calls kmem_cache_create and relies on implicit merging. This uncovered a bug in the slub subsystem - if we delete a cache and immediatelly create another cache with the same attributes, it fails because of duplicate filename in /sys/kernel/slab/. The slub subsystem offloads freeing the cache to a workqueue - and if we create the new cache before the workqueue runs, it complains because of duplicate filename in sysfs. This patch fixes the bug by moving the call of kobject_del from sysfs_slab_remove_workfn to shutdown_cache. kobject_del must be called while we hold slab_mutex - so that the sysfs entry is deleted before a cache with the same attributes could be created. Running device-mapper-test-suite with: dmtest run --suite thin-provisioning -n /commit_failure_causes_fallback/ triggered: Buffer I/O error on dev dm-0, logical block 1572848, async page read device-mapper: thin: 253:1: metadata operation 'dm_pool_alloc_data_block' failed: error = -5 device-mapper: thin: 253:1: aborting current metadata transaction sysfs: cannot create duplicate filename '/kernel/slab/:a-0000144' CPU: 2 PID: 1037 Comm: kworker/u48:1 Not tainted 4.17.0.snitm+ #25 Hardware name: Supermicro SYS-1029P-WTR/X11DDW-L, BIOS 2.0a 12/06/2017 Workqueue: dm-thin do_worker [dm_thin_pool] Call Trace: dump_stack+0x5a/0x73 sysfs_warn_dup+0x58/0x70 sysfs_create_dir_ns+0x77/0x80 kobject_add_internal+0xba/0x2e0 kobject_init_and_add+0x70/0xb0 sysfs_slab_add+0xb1/0x250 __kmem_cache_create+0x116/0x150 create_cache+0xd9/0x1f0 kmem_cache_create_usercopy+0x1c1/0x250 kmem_cache_create+0x18/0x20 dm_bufio_client_create+0x1ae/0x410 [dm_bufio] dm_block_manager_create+0x5e/0x90 [dm_persistent_data] __create_persistent_data_objects+0x38/0x940 [dm_thin_pool] dm_pool_abort_metadata+0x64/0x90 [dm_thin_pool] metadata_operation_failed+0x59/0x100 [dm_thin_pool] alloc_data_block.isra.53+0x86/0x180 [dm_thin_pool] process_cell+0x2a3/0x550 [dm_thin_pool] do_worker+0x28d/0x8f0 [dm_thin_pool] process_one_work+0x171/0x370 worker_thread+0x49/0x3f0 kthread+0xf8/0x130 ret_from_fork+0x35/0x40 kobject_add_internal failed for :a-0000144 with -EEXIST, don't try to register things with the same name in the same directory. kmem_cache_create(dm_bufio_buffer-16) failed with error -17 Link: http://lkml.kernel.org/r/alpine.LRH.2.02.1806151817130.6333@file01.intranet.prod.int.rdu2.redhat.com Signed-off-by: Mikulas Patocka <mpatocka@redhat.com> Reported-by: Mike Snitzer <snitzer@redhat.com> Tested-by: Mike Snitzer <snitzer@redhat.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-06-28 09:26:09 +03:00
void sysfs_slab_unlink(struct kmem_cache *);
void sysfs_slab_release(struct kmem_cache *);
slub: use sysfs'es release mechanism for kmem_cache debugobjects warning during netfilter exit: ------------[ cut here ]------------ WARNING: CPU: 6 PID: 4178 at lib/debugobjects.c:260 debug_print_object+0x8d/0xb0() ODEBUG: free active (active state 0) object type: timer_list hint: delayed_work_timer_fn+0x0/0x20 Modules linked in: CPU: 6 PID: 4178 Comm: kworker/u16:2 Tainted: G W 3.11.0-next-20130906-sasha #3984 Workqueue: netns cleanup_net Call Trace: dump_stack+0x52/0x87 warn_slowpath_common+0x8c/0xc0 warn_slowpath_fmt+0x46/0x50 debug_print_object+0x8d/0xb0 __debug_check_no_obj_freed+0xa5/0x220 debug_check_no_obj_freed+0x15/0x20 kmem_cache_free+0x197/0x340 kmem_cache_destroy+0x86/0xe0 nf_conntrack_cleanup_net_list+0x131/0x170 nf_conntrack_pernet_exit+0x5d/0x70 ops_exit_list+0x5e/0x70 cleanup_net+0xfb/0x1c0 process_one_work+0x338/0x550 worker_thread+0x215/0x350 kthread+0xe7/0xf0 ret_from_fork+0x7c/0xb0 Also during dcookie cleanup: WARNING: CPU: 12 PID: 9725 at lib/debugobjects.c:260 debug_print_object+0x8c/0xb0() ODEBUG: free active (active state 0) object type: timer_list hint: delayed_work_timer_fn+0x0/0x20 Modules linked in: CPU: 12 PID: 9725 Comm: trinity-c141 Not tainted 3.15.0-rc2-next-20140423-sasha-00018-gc4ff6c4 #408 Call Trace: dump_stack (lib/dump_stack.c:52) warn_slowpath_common (kernel/panic.c:430) warn_slowpath_fmt (kernel/panic.c:445) debug_print_object (lib/debugobjects.c:262) __debug_check_no_obj_freed (lib/debugobjects.c:697) debug_check_no_obj_freed (lib/debugobjects.c:726) kmem_cache_free (mm/slub.c:2689 mm/slub.c:2717) kmem_cache_destroy (mm/slab_common.c:363) dcookie_unregister (fs/dcookies.c:302 fs/dcookies.c:343) event_buffer_release (arch/x86/oprofile/../../../drivers/oprofile/event_buffer.c:153) __fput (fs/file_table.c:217) ____fput (fs/file_table.c:253) task_work_run (kernel/task_work.c:125 (discriminator 1)) do_notify_resume (include/linux/tracehook.h:196 arch/x86/kernel/signal.c:751) int_signal (arch/x86/kernel/entry_64.S:807) Sysfs has a release mechanism. Use that to release the kmem_cache structure if CONFIG_SYSFS is enabled. Only slub is changed - slab currently only supports /proc/slabinfo and not /sys/kernel/slab/*. We talked about adding that and someone was working on it. [akpm@linux-foundation.org: fix CONFIG_SYSFS=n build] [akpm@linux-foundation.org: fix CONFIG_SYSFS=n build even more] Signed-off-by: Christoph Lameter <cl@linux.com> Reported-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Acked-by: Greg KH <greg@kroah.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Pekka Enberg <penberg@kernel.org> Cc: Russell King <rmk@arm.linux.org.uk> Cc: Bart Van Assche <bvanassche@acm.org> Cc: Al Viro <viro@ZenIV.linux.org.uk> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-05-06 23:50:08 +04:00
#else
slub: fix failure when we delete and create a slab cache In kernel 4.17 I removed some code from dm-bufio that did slab cache merging (commit 21bb13276768: "dm bufio: remove code that merges slab caches") - both slab and slub support merging caches with identical attributes, so dm-bufio now just calls kmem_cache_create and relies on implicit merging. This uncovered a bug in the slub subsystem - if we delete a cache and immediatelly create another cache with the same attributes, it fails because of duplicate filename in /sys/kernel/slab/. The slub subsystem offloads freeing the cache to a workqueue - and if we create the new cache before the workqueue runs, it complains because of duplicate filename in sysfs. This patch fixes the bug by moving the call of kobject_del from sysfs_slab_remove_workfn to shutdown_cache. kobject_del must be called while we hold slab_mutex - so that the sysfs entry is deleted before a cache with the same attributes could be created. Running device-mapper-test-suite with: dmtest run --suite thin-provisioning -n /commit_failure_causes_fallback/ triggered: Buffer I/O error on dev dm-0, logical block 1572848, async page read device-mapper: thin: 253:1: metadata operation 'dm_pool_alloc_data_block' failed: error = -5 device-mapper: thin: 253:1: aborting current metadata transaction sysfs: cannot create duplicate filename '/kernel/slab/:a-0000144' CPU: 2 PID: 1037 Comm: kworker/u48:1 Not tainted 4.17.0.snitm+ #25 Hardware name: Supermicro SYS-1029P-WTR/X11DDW-L, BIOS 2.0a 12/06/2017 Workqueue: dm-thin do_worker [dm_thin_pool] Call Trace: dump_stack+0x5a/0x73 sysfs_warn_dup+0x58/0x70 sysfs_create_dir_ns+0x77/0x80 kobject_add_internal+0xba/0x2e0 kobject_init_and_add+0x70/0xb0 sysfs_slab_add+0xb1/0x250 __kmem_cache_create+0x116/0x150 create_cache+0xd9/0x1f0 kmem_cache_create_usercopy+0x1c1/0x250 kmem_cache_create+0x18/0x20 dm_bufio_client_create+0x1ae/0x410 [dm_bufio] dm_block_manager_create+0x5e/0x90 [dm_persistent_data] __create_persistent_data_objects+0x38/0x940 [dm_thin_pool] dm_pool_abort_metadata+0x64/0x90 [dm_thin_pool] metadata_operation_failed+0x59/0x100 [dm_thin_pool] alloc_data_block.isra.53+0x86/0x180 [dm_thin_pool] process_cell+0x2a3/0x550 [dm_thin_pool] do_worker+0x28d/0x8f0 [dm_thin_pool] process_one_work+0x171/0x370 worker_thread+0x49/0x3f0 kthread+0xf8/0x130 ret_from_fork+0x35/0x40 kobject_add_internal failed for :a-0000144 with -EEXIST, don't try to register things with the same name in the same directory. kmem_cache_create(dm_bufio_buffer-16) failed with error -17 Link: http://lkml.kernel.org/r/alpine.LRH.2.02.1806151817130.6333@file01.intranet.prod.int.rdu2.redhat.com Signed-off-by: Mikulas Patocka <mpatocka@redhat.com> Reported-by: Mike Snitzer <snitzer@redhat.com> Tested-by: Mike Snitzer <snitzer@redhat.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-06-28 09:26:09 +03:00
static inline void sysfs_slab_unlink(struct kmem_cache *s)
{
}
static inline void sysfs_slab_release(struct kmem_cache *s)
slub: use sysfs'es release mechanism for kmem_cache debugobjects warning during netfilter exit: ------------[ cut here ]------------ WARNING: CPU: 6 PID: 4178 at lib/debugobjects.c:260 debug_print_object+0x8d/0xb0() ODEBUG: free active (active state 0) object type: timer_list hint: delayed_work_timer_fn+0x0/0x20 Modules linked in: CPU: 6 PID: 4178 Comm: kworker/u16:2 Tainted: G W 3.11.0-next-20130906-sasha #3984 Workqueue: netns cleanup_net Call Trace: dump_stack+0x52/0x87 warn_slowpath_common+0x8c/0xc0 warn_slowpath_fmt+0x46/0x50 debug_print_object+0x8d/0xb0 __debug_check_no_obj_freed+0xa5/0x220 debug_check_no_obj_freed+0x15/0x20 kmem_cache_free+0x197/0x340 kmem_cache_destroy+0x86/0xe0 nf_conntrack_cleanup_net_list+0x131/0x170 nf_conntrack_pernet_exit+0x5d/0x70 ops_exit_list+0x5e/0x70 cleanup_net+0xfb/0x1c0 process_one_work+0x338/0x550 worker_thread+0x215/0x350 kthread+0xe7/0xf0 ret_from_fork+0x7c/0xb0 Also during dcookie cleanup: WARNING: CPU: 12 PID: 9725 at lib/debugobjects.c:260 debug_print_object+0x8c/0xb0() ODEBUG: free active (active state 0) object type: timer_list hint: delayed_work_timer_fn+0x0/0x20 Modules linked in: CPU: 12 PID: 9725 Comm: trinity-c141 Not tainted 3.15.0-rc2-next-20140423-sasha-00018-gc4ff6c4 #408 Call Trace: dump_stack (lib/dump_stack.c:52) warn_slowpath_common (kernel/panic.c:430) warn_slowpath_fmt (kernel/panic.c:445) debug_print_object (lib/debugobjects.c:262) __debug_check_no_obj_freed (lib/debugobjects.c:697) debug_check_no_obj_freed (lib/debugobjects.c:726) kmem_cache_free (mm/slub.c:2689 mm/slub.c:2717) kmem_cache_destroy (mm/slab_common.c:363) dcookie_unregister (fs/dcookies.c:302 fs/dcookies.c:343) event_buffer_release (arch/x86/oprofile/../../../drivers/oprofile/event_buffer.c:153) __fput (fs/file_table.c:217) ____fput (fs/file_table.c:253) task_work_run (kernel/task_work.c:125 (discriminator 1)) do_notify_resume (include/linux/tracehook.h:196 arch/x86/kernel/signal.c:751) int_signal (arch/x86/kernel/entry_64.S:807) Sysfs has a release mechanism. Use that to release the kmem_cache structure if CONFIG_SYSFS is enabled. Only slub is changed - slab currently only supports /proc/slabinfo and not /sys/kernel/slab/*. We talked about adding that and someone was working on it. [akpm@linux-foundation.org: fix CONFIG_SYSFS=n build] [akpm@linux-foundation.org: fix CONFIG_SYSFS=n build even more] Signed-off-by: Christoph Lameter <cl@linux.com> Reported-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Acked-by: Greg KH <greg@kroah.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Pekka Enberg <penberg@kernel.org> Cc: Russell King <rmk@arm.linux.org.uk> Cc: Bart Van Assche <bvanassche@acm.org> Cc: Al Viro <viro@ZenIV.linux.org.uk> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-05-06 23:50:08 +04:00
{
}
#endif
void object_err(struct kmem_cache *s, struct page *page,
u8 *object, char *reason);
void *fixup_red_left(struct kmem_cache *s, void *p);
static inline void *nearest_obj(struct kmem_cache *cache, struct page *page,
void *x) {
void *object = x - (x - page_address(page)) % cache->size;
void *last_object = page_address(page) +
(page->objects - 1) * cache->size;
void *result = (unlikely(object > last_object)) ? last_object : object;
result = fixup_red_left(cache, result);
return result;
}
SLUB core This is a new slab allocator which was motivated by the complexity of the existing code in mm/slab.c. It attempts to address a variety of concerns with the existing implementation. A. Management of object queues A particular concern was the complex management of the numerous object queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for each allocating CPU and use objects from a slab directly instead of queueing them up. B. Storage overhead of object queues SLAB Object queues exist per node, per CPU. The alien cache queue even has a queue array that contain a queue for each processor on each node. For very large systems the number of queues and the number of objects that may be caught in those queues grows exponentially. On our systems with 1k nodes / processors we have several gigabytes just tied up for storing references to objects for those queues This does not include the objects that could be on those queues. One fears that the whole memory of the machine could one day be consumed by those queues. C. SLAB meta data overhead SLAB has overhead at the beginning of each slab. This means that data cannot be naturally aligned at the beginning of a slab block. SLUB keeps all meta data in the corresponding page_struct. Objects can be naturally aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte boundaries and can fit tightly into a 4k page with no bytes left over. SLAB cannot do this. D. SLAB has a complex cache reaper SLUB does not need a cache reaper for UP systems. On SMP systems the per CPU slab may be pushed back into partial list but that operation is simple and does not require an iteration over a list of objects. SLAB expires per CPU, shared and alien object queues during cache reaping which may cause strange hold offs. E. SLAB has complex NUMA policy layer support SLUB pushes NUMA policy handling into the page allocator. This means that allocation is coarser (SLUB does interleave on a page level) but that situation was also present before 2.6.13. SLABs application of policies to individual slab objects allocated in SLAB is certainly a performance concern due to the frequent references to memory policies which may lead a sequence of objects to come from one node after another. SLUB will get a slab full of objects from one node and then will switch to the next. F. Reduction of the size of partial slab lists SLAB has per node partial lists. This means that over time a large number of partial slabs may accumulate on those lists. These can only be reused if allocator occur on specific nodes. SLUB has a global pool of partial slabs and will consume slabs from that pool to decrease fragmentation. G. Tunables SLAB has sophisticated tuning abilities for each slab cache. One can manipulate the queue sizes in detail. However, filling the queues still requires the uses of the spin lock to check out slabs. SLUB has a global parameter (min_slab_order) for tuning. Increasing the minimum slab order can decrease the locking overhead. The bigger the slab order the less motions of pages between per CPU and partial lists occur and the better SLUB will be scaling. G. Slab merging We often have slab caches with similar parameters. SLUB detects those on boot up and merges them into the corresponding general caches. This leads to more effective memory use. About 50% of all caches can be eliminated through slab merging. This will also decrease slab fragmentation because partial allocated slabs can be filled up again. Slab merging can be switched off by specifying slub_nomerge on boot up. Note that merging can expose heretofore unknown bugs in the kernel because corrupted objects may now be placed differently and corrupt differing neighboring objects. Enable sanity checks to find those. H. Diagnostics The current slab diagnostics are difficult to use and require a recompilation of the kernel. SLUB contains debugging code that is always available (but is kept out of the hot code paths). SLUB diagnostics can be enabled via the "slab_debug" option. Parameters can be specified to select a single or a group of slab caches for diagnostics. This means that the system is running with the usual performance and it is much more likely that race conditions can be reproduced. I. Resiliency If basic sanity checks are on then SLUB is capable of detecting common error conditions and recover as best as possible to allow the system to continue. J. Tracing Tracing can be enabled via the slab_debug=T,<slabcache> option during boot. SLUB will then protocol all actions on that slabcache and dump the object contents on free. K. On demand DMA cache creation. Generally DMA caches are not needed. If a kmalloc is used with __GFP_DMA then just create this single slabcache that is needed. For systems that have no ZONE_DMA requirement the support is completely eliminated. L. Performance increase Some benchmarks have shown speed improvements on kernbench in the range of 5-10%. The locking overhead of slub is based on the underlying base allocation size. If we can reliably allocate larger order pages then it is possible to increase slub performance much further. The anti-fragmentation patches may enable further performance increases. Tested on: i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator SLUB Boot options slub_nomerge Disable merging of slabs slub_min_order=x Require a minimum order for slab caches. This increases the managed chunk size and therefore reduces meta data and locking overhead. slub_min_objects=x Mininum objects per slab. Default is 8. slub_max_order=x Avoid generating slabs larger than order specified. slub_debug Enable all diagnostics for all caches slub_debug=<options> Enable selective options for all caches slub_debug=<o>,<cache> Enable selective options for a certain set of caches Available Debug options F Double Free checking, sanity and resiliency R Red zoning P Object / padding poisoning U Track last free / alloc T Trace all allocs / frees (only use for individual slabs). To use SLUB: Apply this patch and then select SLUB as the default slab allocator. [hugh@veritas.com: fix an oops-causing locking error] [akpm@linux-foundation.org: various stupid cleanups and small fixes] Signed-off-by: Christoph Lameter <clameter@sgi.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-05-07 01:49:36 +04:00
#endif /* _LINUX_SLUB_DEF_H */