328 строки
13 KiB
Plaintext
328 строки
13 KiB
Plaintext
Kernel level exception handling in Linux
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Commentary by Joerg Pommnitz <joerg@raleigh.ibm.com>
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When a process runs in kernel mode, it often has to access user
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mode memory whose address has been passed by an untrusted program.
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To protect itself the kernel has to verify this address.
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In older versions of Linux this was done with the
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int verify_area(int type, const void * addr, unsigned long size)
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function (which has since been replaced by access_ok()).
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This function verified that the memory area starting at address
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'addr' and of size 'size' was accessible for the operation specified
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in type (read or write). To do this, verify_read had to look up the
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virtual memory area (vma) that contained the address addr. In the
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normal case (correctly working program), this test was successful.
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It only failed for a few buggy programs. In some kernel profiling
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tests, this normally unneeded verification used up a considerable
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amount of time.
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To overcome this situation, Linus decided to let the virtual memory
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hardware present in every Linux-capable CPU handle this test.
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How does this work?
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Whenever the kernel tries to access an address that is currently not
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accessible, the CPU generates a page fault exception and calls the
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page fault handler
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void do_page_fault(struct pt_regs *regs, unsigned long error_code)
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in arch/x86/mm/fault.c. The parameters on the stack are set up by
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the low level assembly glue in arch/x86/kernel/entry_32.S. The parameter
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regs is a pointer to the saved registers on the stack, error_code
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contains a reason code for the exception.
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do_page_fault first obtains the unaccessible address from the CPU
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control register CR2. If the address is within the virtual address
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space of the process, the fault probably occurred, because the page
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was not swapped in, write protected or something similar. However,
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we are interested in the other case: the address is not valid, there
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is no vma that contains this address. In this case, the kernel jumps
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to the bad_area label.
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There it uses the address of the instruction that caused the exception
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(i.e. regs->eip) to find an address where the execution can continue
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(fixup). If this search is successful, the fault handler modifies the
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return address (again regs->eip) and returns. The execution will
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continue at the address in fixup.
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Where does fixup point to?
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Since we jump to the contents of fixup, fixup obviously points
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to executable code. This code is hidden inside the user access macros.
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I have picked the get_user macro defined in arch/x86/include/asm/uaccess.h
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as an example. The definition is somewhat hard to follow, so let's peek at
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the code generated by the preprocessor and the compiler. I selected
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the get_user call in drivers/char/sysrq.c for a detailed examination.
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The original code in sysrq.c line 587:
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get_user(c, buf);
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The preprocessor output (edited to become somewhat readable):
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(
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{
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long __gu_err = - 14 , __gu_val = 0;
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const __typeof__(*( ( buf ) )) *__gu_addr = ((buf));
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if (((((0 + current_set[0])->tss.segment) == 0x18 ) ||
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(((sizeof(*(buf))) <= 0xC0000000UL) &&
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((unsigned long)(__gu_addr ) <= 0xC0000000UL - (sizeof(*(buf)))))))
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do {
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__gu_err = 0;
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switch ((sizeof(*(buf)))) {
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case 1:
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__asm__ __volatile__(
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"1: mov" "b" " %2,%" "b" "1\n"
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"2:\n"
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".section .fixup,\"ax\"\n"
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"3: movl %3,%0\n"
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" xor" "b" " %" "b" "1,%" "b" "1\n"
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" jmp 2b\n"
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".section __ex_table,\"a\"\n"
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" .align 4\n"
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" .long 1b,3b\n"
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".text" : "=r"(__gu_err), "=q" (__gu_val): "m"((*(struct __large_struct *)
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( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err )) ;
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break;
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case 2:
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__asm__ __volatile__(
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"1: mov" "w" " %2,%" "w" "1\n"
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"2:\n"
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".section .fixup,\"ax\"\n"
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"3: movl %3,%0\n"
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" xor" "w" " %" "w" "1,%" "w" "1\n"
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" jmp 2b\n"
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".section __ex_table,\"a\"\n"
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" .align 4\n"
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" .long 1b,3b\n"
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".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
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( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err ));
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break;
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case 4:
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__asm__ __volatile__(
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"1: mov" "l" " %2,%" "" "1\n"
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"2:\n"
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".section .fixup,\"ax\"\n"
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"3: movl %3,%0\n"
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" xor" "l" " %" "" "1,%" "" "1\n"
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" jmp 2b\n"
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".section __ex_table,\"a\"\n"
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" .align 4\n" " .long 1b,3b\n"
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".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
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( __gu_addr )) ), "i"(- 14 ), "0"(__gu_err));
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break;
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default:
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(__gu_val) = __get_user_bad();
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}
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} while (0) ;
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((c)) = (__typeof__(*((buf))))__gu_val;
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__gu_err;
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}
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);
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WOW! Black GCC/assembly magic. This is impossible to follow, so let's
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see what code gcc generates:
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> xorl %edx,%edx
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> movl current_set,%eax
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> cmpl $24,788(%eax)
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> je .L1424
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> cmpl $-1073741825,64(%esp)
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> ja .L1423
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> .L1424:
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> movl %edx,%eax
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> movl 64(%esp),%ebx
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> #APP
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> 1: movb (%ebx),%dl /* this is the actual user access */
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> 2:
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> .section .fixup,"ax"
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> 3: movl $-14,%eax
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> xorb %dl,%dl
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> jmp 2b
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> .section __ex_table,"a"
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> .align 4
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> .long 1b,3b
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> .text
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> #NO_APP
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> .L1423:
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> movzbl %dl,%esi
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The optimizer does a good job and gives us something we can actually
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understand. Can we? The actual user access is quite obvious. Thanks
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to the unified address space we can just access the address in user
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memory. But what does the .section stuff do?????
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To understand this we have to look at the final kernel:
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> objdump --section-headers vmlinux
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>
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> vmlinux: file format elf32-i386
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>
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> Sections:
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> Idx Name Size VMA LMA File off Algn
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> 0 .text 00098f40 c0100000 c0100000 00001000 2**4
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> CONTENTS, ALLOC, LOAD, READONLY, CODE
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> 1 .fixup 000016bc c0198f40 c0198f40 00099f40 2**0
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> CONTENTS, ALLOC, LOAD, READONLY, CODE
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> 2 .rodata 0000f127 c019a5fc c019a5fc 0009b5fc 2**2
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> CONTENTS, ALLOC, LOAD, READONLY, DATA
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> 3 __ex_table 000015c0 c01a9724 c01a9724 000aa724 2**2
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> CONTENTS, ALLOC, LOAD, READONLY, DATA
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> 4 .data 0000ea58 c01abcf0 c01abcf0 000abcf0 2**4
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> CONTENTS, ALLOC, LOAD, DATA
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> 5 .bss 00018e21 c01ba748 c01ba748 000ba748 2**2
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> ALLOC
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> 6 .comment 00000ec4 00000000 00000000 000ba748 2**0
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> CONTENTS, READONLY
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> 7 .note 00001068 00000ec4 00000ec4 000bb60c 2**0
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> CONTENTS, READONLY
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There are obviously 2 non standard ELF sections in the generated object
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file. But first we want to find out what happened to our code in the
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final kernel executable:
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> objdump --disassemble --section=.text vmlinux
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>
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> c017e785 <do_con_write+c1> xorl %edx,%edx
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> c017e787 <do_con_write+c3> movl 0xc01c7bec,%eax
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> c017e78c <do_con_write+c8> cmpl $0x18,0x314(%eax)
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> c017e793 <do_con_write+cf> je c017e79f <do_con_write+db>
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> c017e795 <do_con_write+d1> cmpl $0xbfffffff,0x40(%esp,1)
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> c017e79d <do_con_write+d9> ja c017e7a7 <do_con_write+e3>
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> c017e79f <do_con_write+db> movl %edx,%eax
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> c017e7a1 <do_con_write+dd> movl 0x40(%esp,1),%ebx
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> c017e7a5 <do_con_write+e1> movb (%ebx),%dl
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> c017e7a7 <do_con_write+e3> movzbl %dl,%esi
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The whole user memory access is reduced to 10 x86 machine instructions.
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The instructions bracketed in the .section directives are no longer
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in the normal execution path. They are located in a different section
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of the executable file:
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> objdump --disassemble --section=.fixup vmlinux
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>
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> c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
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> c0199ffa <.fixup+10ba> xorb %dl,%dl
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> c0199ffc <.fixup+10bc> jmp c017e7a7 <do_con_write+e3>
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And finally:
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> objdump --full-contents --section=__ex_table vmlinux
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>
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> c01aa7c4 93c017c0 e09f19c0 97c017c0 99c017c0 ................
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> c01aa7d4 f6c217c0 e99f19c0 a5e717c0 f59f19c0 ................
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> c01aa7e4 080a18c0 01a019c0 0a0a18c0 04a019c0 ................
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or in human readable byte order:
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> c01aa7c4 c017c093 c0199fe0 c017c097 c017c099 ................
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> c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
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^^^^^^^^^^^^^^^^^
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this is the interesting part!
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> c01aa7e4 c0180a08 c019a001 c0180a0a c019a004 ................
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What happened? The assembly directives
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.section .fixup,"ax"
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.section __ex_table,"a"
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told the assembler to move the following code to the specified
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sections in the ELF object file. So the instructions
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3: movl $-14,%eax
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xorb %dl,%dl
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jmp 2b
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ended up in the .fixup section of the object file and the addresses
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.long 1b,3b
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ended up in the __ex_table section of the object file. 1b and 3b
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are local labels. The local label 1b (1b stands for next label 1
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backward) is the address of the instruction that might fault, i.e.
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in our case the address of the label 1 is c017e7a5:
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the original assembly code: > 1: movb (%ebx),%dl
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and linked in vmlinux : > c017e7a5 <do_con_write+e1> movb (%ebx),%dl
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The local label 3 (backwards again) is the address of the code to handle
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the fault, in our case the actual value is c0199ff5:
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the original assembly code: > 3: movl $-14,%eax
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and linked in vmlinux : > c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
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The assembly code
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> .section __ex_table,"a"
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> .align 4
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> .long 1b,3b
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becomes the value pair
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> c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
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^this is ^this is
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1b 3b
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c017e7a5,c0199ff5 in the exception table of the kernel.
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So, what actually happens if a fault from kernel mode with no suitable
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vma occurs?
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1.) access to invalid address:
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> c017e7a5 <do_con_write+e1> movb (%ebx),%dl
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2.) MMU generates exception
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3.) CPU calls do_page_fault
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4.) do page fault calls search_exception_table (regs->eip == c017e7a5);
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5.) search_exception_table looks up the address c017e7a5 in the
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exception table (i.e. the contents of the ELF section __ex_table)
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and returns the address of the associated fault handle code c0199ff5.
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6.) do_page_fault modifies its own return address to point to the fault
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handle code and returns.
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7.) execution continues in the fault handling code.
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8.) 8a) EAX becomes -EFAULT (== -14)
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8b) DL becomes zero (the value we "read" from user space)
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8c) execution continues at local label 2 (address of the
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instruction immediately after the faulting user access).
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The steps 8a to 8c in a certain way emulate the faulting instruction.
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That's it, mostly. If you look at our example, you might ask why
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we set EAX to -EFAULT in the exception handler code. Well, the
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get_user macro actually returns a value: 0, if the user access was
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successful, -EFAULT on failure. Our original code did not test this
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return value, however the inline assembly code in get_user tries to
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return -EFAULT. GCC selected EAX to return this value.
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NOTE:
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Due to the way that the exception table is built and needs to be ordered,
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only use exceptions for code in the .text section. Any other section
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will cause the exception table to not be sorted correctly, and the
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exceptions will fail.
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Things changed when 64-bit support was added to x86 Linux. Rather than
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double the size of the exception table by expanding the two entries
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from 32-bits to 64 bits, a clever trick was used to store addresses
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as relative offsets from the table itself. The assembly code changed
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from:
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.long 1b,3b
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to:
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.long (from) - .
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.long (to) - .
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and the C-code that uses these values converts back to absolute addresses
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like this:
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ex_insn_addr(const struct exception_table_entry *x)
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{
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return (unsigned long)&x->insn + x->insn;
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}
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In v4.6 the exception table entry was expanded with a new field "handler".
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This is also 32-bits wide and contains a third relative function
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pointer which points to one of:
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1) int ex_handler_default(const struct exception_table_entry *fixup)
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This is legacy case that just jumps to the fixup code
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2) int ex_handler_fault(const struct exception_table_entry *fixup)
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This case provides the fault number of the trap that occurred at
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entry->insn. It is used to distinguish page faults from machine
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check.
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3) int ex_handler_ext(const struct exception_table_entry *fixup)
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This case is used for uaccess_err ... we need to set a flag
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in the task structure. Before the handler functions existed this
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case was handled by adding a large offset to the fixup to tag
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it as special.
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More functions can easily be added.
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