linux下糟糕的异常处理方式
时间：2007-02-17 来源：PHP爱好者
linux下发生异常，芯片会自动产生一个异常中断。在这异常中断处理程序中会判断异常来自用户程序或者内核，如果是发生在用户程序，那么会产生一个异常信号，再根据异常信号的回调函数通知用户程序发生异常。如果发生在内核里面，那么就会搜索内核模块的异常结构表，找到相应的处理调用地址，修改异常中断的返回地址为异常处理的地址，中断返回的时候程序就跳到异常处理程序处理执行了。但具体这两种处理方法都很糟糕，下面简要分析一下。
linux系统把所有进程数据结构都放于内核，这就增加了一些不必要的切换时间。 linux可以通过系统调用，安装信号的回调函数，这回调函数指针存放在内核的进程数据结构里面。这点windows处理得比较好，windows把进程数据结构分成了两部分，一部分敏感数据放于内核的进程数据结构里面，加以保护，另一部分不敏感数据就放于用户空间，这样当访问那些不加保护的数据时，就不用切换到内核，节约了时间。像windows下异常处理，也是一种回调函数，但因为结构放于用户空间，安装的时候就很方便，也节约切换时间。
上面那一点只是效率问题，但linux内核的异常处理那才是糟糕。先介绍一下linux内核的异常处理结构吧，看明白了你自然就知道糟糕到什么程度了。要了解这，显然应该是先从异常中断入手。下面主要是x86芯片的一些处理，但别的芯片下的也应该差不多。
文件：entry.S:
ENTRY(general_protection)
pushl $ SYMBOL_NAME(do_general_protection)
jmp error_code
这是异常中断入口，显然会执行do_general_protection。
文件traps.c:
asmlinkage void do_general_protection(struct pt_regs * regs, long error_code)
{
if (regs->eflags & VM_MASK)
goto gp_in_vm86;
/*
虚拟8086下发生的异常否
*/
if (!(regs->xcs & 3))
goto gp_in_kernel;
/*
内核发生的异常否
*/
current->tss.error_code = error_code;
current->tss.trap_no = 13;
force_sig(SIGSEGV, current);
/*
用户程序发生的异常，产生异常信号，
根据异常信号的句柄回调处理函数
*/
return;
gp_in_vm86:
lock_kernel();
handle_vm86_fault((struct kernel_vm86_regs *) regs, error_code);
/*
虚拟8086的处理
*/
unlock_kernel();
return;
gp_in_kernel:
{
unsigned long fixup;
fixup = search_exception_table(regs->eip);
/*
根据异常时的eip搜索异常结构链
找到处理程序地址
*/
if (fixup) {
regs->eip = fixup;
/*
找到异常处理地址，修改中断返回地址，中断返回时跳到异常处理程序处
*/
return;
}
die("general protection fault", regs, error_code);
/*
没找到异常处理程序地址，显示内核异常信息后死机
*/
}
}
搜索异常处理程序代码文件extable.c:
extern const struct exception_table_entry __start___ex_table[];
extern const struct exception_table_entry __stop___ex_table[];
unsigned long search_exception_table(unsigned long addr)
{
unsigned long ret;
#ifndef CONFIG_MODULES
/* There is only the kernel to search. */
ret = search_one_table(__start___ex_table, __stop___ex_table-1, addr);
if (ret) return ret;
#else
/* The kernel is the last "module" -- no need to treat it special. */
struct module *mp;
for (mp = module_list; mp != NULL; mp = mp->next) {
if (mp->ex_table_start == NULL)
continue;
ret = search_one_table(mp->ex_table_start,
mp->ex_table_end - 1, addr);
if (ret) return ret;
}
#endif
return 0;
}
static inline unsigned long
search_one_table(const struct exception_table_entry *first,
const struct exception_table_entry *last,
unsigned long value)
{
while (first <= last) {
const struct exception_table_entry *mid;
long diff;
mid = (last - first) / 2 + first;
diff = mid->insn - value;
if (diff == 0)
return mid->fixup;
else if (diff <0)
first = mid+1;
else
last = mid-1;
}
return 0;
}
看看上面搜索异常处理程序的算法就知道了，有个异常模块表，保存会发生异常时候的eip和异常处理程序指针，发生异常的时候就根据异常时候的eip搜索表里面的eip，发现相等就找到了异常处理指针。这是什么意思呢，就是说你编写的内核程序必须精确的知道哪条指令可能会发生异常，要求真够高的。想想windows下的异常编程是多么轻松？程序员只需要知道哪一段程序可能出现异常，就只需要一个括号一个异常语句保护这段程序就是了。
光看上面算法可能对其这异常的处理还没怎么有感性认识，那么我们再看看其内核的异常形式、编写方式吧。
下面我们再看看exception.txt的一些说明:
查看内核模块表：
> Sections:
> Idx Name Size VMA LMA File off Algn
> 0 .text 00098f40 c0100000 c0100000 00001000 2**4
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 1 .fixup 000016bc c0198f40 c0198f40 00099f40 2**0
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 2 .rodata 0000f127 c019a5fc c019a5fc 0009b5fc 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 3 __ex_table 000015c0 c01a9724 c01a9724 000aa724 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 4 .data 0000ea58 c01abcf0 c01abcf0 000abcf0 2**4
> CONTENTS, ALLOC, LOAD, DATA
> 5 .bss 00018e21 c01ba748 c01ba748 000ba748 2**2
> ALLOC
> 6 .comment 00000ec4 00000000 00000000 000ba748 2**0
> CONTENTS, READONLY
> 7 .note 00001068 00000ec4 00000ec4 000bb60c 2**0
> CONTENTS, READONLY
看模块__ex_table就是会出现异常的一些程序eip，fixup就是相应的异常处理程序地址。这显然这异常结构是静态的，与windows动态的链表形式有很大的分别。
再看看异常程序的编写，下面是get_user(c, buf)的一段代码：
switch ((sizeof(*(buf)))) {
case 1:
__asm__ __volatile__(
"1: mov" "b" " %2,%" "b" "1n"
/*
这语句可能发生异常
*/
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
/*
异常处理程序
*/
" xor" "b" " %" "b" "1,%" "b" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n"
" .long 1b,3bn"
/*
1b,3b就是对应1：，3：就是可能会发生异常的eip与异常后的处理程序
*/
".text" : "=r"(__gu_err), "=q" (__gu_val): "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err )) ;
break;
看看上面代码，是不是要严格的知道哪条指令可能会产生异常？你也清楚了linux整个异常处理方式了吧。编写是不是也太麻烦？弄不好你不知道哪条语句会发生异常，而执行到那发生了异常，那么系统就会出现可怕的异常提示后死机了。
这点完全可以用一个链表的方式处理，哪段程序可能发生异常，就安装异常处理程序指针到链表里面，执行完这段代码就从链表里面删除这段异常处理代码指针，如果这段代码发生异常，系统的异常中断处理程序只需要调用这链表的异常处理程序指针就是了，不要匹配发生异常时的eip。linux的处理主要是因为处理成了一个静态的链表，而不是动态的，这就使得需要确定发生异常时的异常处理指针，所以就增加了一个检测eip完全相等的匹配条件，而这就造成了程序编写上的苦难。这处理方式不用我说你就会知道是多么的糟糕了，弄不好就会留下一些会让系统崩溃的地方呢。
附exception.txt:
Kernel level exception handling in Linux 2.1.8
Commentary by Joerg Pommnitz
When a process runs in kernel mode, it often has to access user
mode memory whose address has been passed by an untrusted program.
To protect itself the kernel has to verify this address.
In older versions of Linux this was done with the
int verify_area(int type, const void * addr, unsigned long size)
function.
This function verified that the memory area starting at address
addr and of size size was accessible for the operation specified
in type (read or write). To do this, verify_read had to look up the
virtual memory area (vma) that contained the address addr. In the
normal case (correctly working program), this test was successful.
It only failed for a few buggy programs. In some kernel profiling
tests, this normally unneeded verification used up a considerable
amount of time.
To overcome this situation, Linus decided to let the virtual memory
hardware present in every Linux-capable CPU handle this test.
How does this work?
Whenever the kernel tries to access an address that is currently not
accessible, the CPU generates a page fault exception and calls the
page fault handler
void do_page_fault(struct pt_regs *regs, unsigned long error_code)
in arch/i386/mm/fault.c. The parameters on the stack are set up by
the low level assembly glue in arch/i386/kernel/entry.S. The parameter
regs is a pointer to the saved registers on the stack, error_code
contains a reason code for the exception.
do_page_fault first obtains the unaccessible address from the CPU
control register CR2. If the address is within the virtual address
space of the process, the fault probably occurred, because the page
was not swapped in, write protected or something similar. However,
we are interested in the other case: the address is not valid, there
is no vma that contains this address. In this case, the kernel jumps
to the bad_area label.
There it uses the address of the instruction that caused the exception
(i.e. regs->eip) to find an address where the execution can continue
(fixup). If this search is successful, the fault handler modifies the
return address (again regs->eip) and returns. The execution will
continue at the address in fixup.
Where does fixup point to?
Since we jump to the contents of fixup, fixup obviously points
to executable code. This code is hidden inside the user access macros.
I have picked the get_user macro defined in include/asm/uaccess.h as an
example. The definition is somewhat hard to follow, so let's peek at
the code generated by the preprocessor and the compiler. I selected
the get_user call in drivers/char/console.c for a detailed examination.
The original code in console.c line 1405:
get_user(c, buf);
The preprocessor output (edited to become somewhat readable):
(
{
long __gu_err = - 14 , __gu_val = 0;
const __typeof__(*( ( buf ) )) *__gu_addr = ((buf));
if (((((0 + current_set[0])->tss.segment) == 0x18 ) ||
(((sizeof(*(buf))) <= 0xC0000000UL) &&
((unsigned long)(__gu_addr ) <= 0xC0000000UL - (sizeof(*(buf)))))))
do {
__gu_err = 0;
switch ((sizeof(*(buf)))) {
case 1:
__asm__ __volatile__(
"1: mov" "b" " %2,%" "b" "1n"
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
" xor" "b" " %" "b" "1,%" "b" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n"
" .long 1b,3bn"
".text" : "=r"(__gu_err), "=q" (__gu_val): "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err )) ;
break;
case 2:
__asm__ __volatile__(
"1: mov" "w" " %2,%" "w" "1n"
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
" xor" "w" " %" "w" "1,%" "w" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n"
" .long 1b,3bn"
".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err ));
break;
case 4:
__asm__ __volatile__(
"1: mov" "l" " %2,%" "" "1n"
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
" xor" "l" " %" "" "1,%" "" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n" " .long 1b,3bn"
".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"(__gu_err));
break;
default:
(__gu_val) = __get_user_bad();
}
} while (0) ;
((c)) = (__typeof__(*((buf))))__gu_val;
__gu_err;
}
);
WOW! Black GCC/assembly magic. This is impossible to follow, so let's
see what code gcc generates:
> xorl %edx,%edx
> movl current_set,%eax
> cmpl $24,788(%eax)
> je .L1424
> cmpl $-1073741825,64(%esp)
> ja .L1423
> .L1424:
> movl %edx,%eax
> movl 64(%esp),%ebx
> #APP
> 1: movb (%ebx),%dl /* this is the actual user access */
> 2:
> .section .fixup,"ax"
> 3: movl $-14,%eax
> xorb %dl,%dl
> jmp 2b
> .section __ex_table,"a"
> .align 4
> .long 1b,3b
> .text
> #NO_APP
> .L1423:
> movzbl %dl,%esi
The optimizer does a good job and gives us something we can actually
understand. Can we? The actual user access is quite obvious. Thanks
to the unified address space we can just access the address in user
memory. But what does the .section stuff do?????
To understand this we have to look at the final kernel:
> objdump --section-headers vmlinux
>
> vmlinux: file format elf32-i386
>
> Sections:
> Idx Name Size VMA LMA File off Algn
> 0 .text 00098f40 c0100000 c0100000 00001000 2**4
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 1 .fixup 000016bc c0198f40 c0198f40 00099f40 2**0
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 2 .rodata 0000f127 c019a5fc c019a5fc 0009b5fc 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 3 __ex_table 000015c0 c01a9724 c01a9724 000aa724 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 4 .data 0000ea58 c01abcf0 c01abcf0 000abcf0 2**4
> CONTENTS, ALLOC, LOAD, DATA
> 5 .bss 00018e21 c01ba748 c01ba748 000ba748 2**2
> ALLOC
> 6 .comment 00000ec4 00000000 00000000 000ba748 2**0
> CONTENTS, READONLY
> 7 .note 00001068 00000ec4 00000ec4 000bb60c 2**0
> CONTENTS, READONLY
There are obviously 2 non standard ELF sections in the generated object
file. But first we want to find out what happened to our code in the
final kernel executable:
> objdump --disassemble --section=.text vmlinux
>
> c017e785 xorl %edx,%edx
> c017e787 movl 0xc01c7bec,%eax
> c017e78c cmpl $0x18,0x314(%eax)
> c017e793 je c017e79f
> c017e795 cmpl $0xbfffffff,0x40(%esp,1)
> c017e79d ja c017e7a7
> c017e79f movl %edx,%eax
> c017e7a1 movl 0x40(%esp,1),%ebx
> c017e7a5 movb (%ebx),%dl
> c017e7a7 movzbl %dl,%esi
The whole user memory access is reduced to 10 x86 machine instructions.
The instructions bracketed in the .section directives are not longer
in the normal execution path. They are located in a different section
of the executable file:
> objdump --disassemble --section=.fixup vmlinux
>
> c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
> c0199ffa <.fixup+10ba> xorb %dl,%dl
> c0199ffc <.fixup+10bc> jmp c017e7a7
And finally:
> objdump --full-contents --section=__ex_table vmlinux
>
> c01aa7c4 93c017c0 e09f19c0 97c017c0 99c017c0 ................
> c01aa7d4 f6c217c0 e99f19c0 a5e717c0 f59f19c0 ................
> c01aa7e4 080a18c0 01a019c0 0a0a18c0 04a019c0 ................
or in human readable byte order:
> c01aa7c4 c017c093 c0199fe0 c017c097 c017c099 ................
> c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
^^^^^^^^^^^^^^^^^
this is the interesting part!
> c01aa7e4 c0180a08 c019a001 c0180a0a c019a004 ................
What happened? The assembly directives
.section .fixup,"ax"
.section __ex_table,"a"
told the assembler to move the following code to the specified
sections in the ELF object file. So the instructions
3: movl $-14,%eax
xorb %dl,%dl
jmp 2b
ended up in the .fixup section of the object file and the addresses
.long 1b,3b
ended up in the __ex_table section of the object file. 1b and 3b
are local labels. The local label 1b (1b stands for next label 1
backward) is the address of the instruction that might fault, i.e.
in our case the address of the label 1 is c017e7a5:
the original assembly code: > 1: movb (%ebx),%dl
and linked in vmlinux : > c017e7a5 movb (%ebx),%dl
The local label 3 (backwards again) is the address of the code to handle
the fault, in our case the actual value is c0199ff5:
the original assembly code: > 3: movl $-14,%eax
and linked in vmlinux : > c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
The assembly code
> .section __ex_table,"a"
> .align 4
> .long 1b,3b
becomes the value pair
> c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
^this is ^this is
1b 3b
c017e7a5,c0199ff5 in the exception table of the kernel.
So, what actually happens if a fault from kernel mode with no suitable
vma occurs?
1.) access to invalid address:
> c017e7a5 movb (%ebx),%dl
2.) MMU generates exception
3.) CPU calls do_page_fault
4.) do page fault calls search_exception_table (regs->eip == c017e7a5);
5.) search_exception_table looks up the address c017e7a5 in the
exception table (i.e. the contents of the ELF section __ex_table)
and returns the address of the associated fault handle code c0199ff5.
6.) do_page_fault modifies its own return address to point to the fault
handle code and returns.
7.) execution continues in the fault handling code.
8.) 8a) EAX becomes -EFAULT (== -14)
8b) DL becomes zero (the value we "read" from user space)
8c) execution continues at local label 2 (address of the
instruction immediately after the faulting user access).
The steps 8a to 8c in a certain way emulate the faulting instruction.
That's it, mostly. If you look at our example, you might ask why
we set EAX to -EFAULT in the exception handler code. Well, the
get_user macro actually returns a value: 0, if the user access was
successful, -EFAULT on failure. Our original code did not test this
return value, however the inline assembly code in get_user tries to
return -EFAULT. GCC selected EAX to return this value.
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linux系统把所有进程数据结构都放于内核，这就增加了一些不必要的切换时间。 linux可以通过系统调用，安装信号的回调函数，这回调函数指针存放在内核的进程数据结构里面。这点windows处理得比较好，windows把进程数据结构分成了两部分，一部分敏感数据放于内核的进程数据结构里面，加以保护，另一部分不敏感数据就放于用户空间，这样当访问那些不加保护的数据时，就不用切换到内核，节约了时间。像windows下异常处理，也是一种回调函数，但因为结构放于用户空间，安装的时候就很方便，也节约切换时间。
上面那一点只是效率问题，但linux内核的异常处理那才是糟糕。先介绍一下linux内核的异常处理结构吧，看明白了你自然就知道糟糕到什么程度了。要了解这，显然应该是先从异常中断入手。下面主要是x86芯片的一些处理，但别的芯片下的也应该差不多。
文件：entry.S:
ENTRY(general_protection)
pushl $ SYMBOL_NAME(do_general_protection)
jmp error_code
这是异常中断入口，显然会执行do_general_protection。
文件traps.c:
asmlinkage void do_general_protection(struct pt_regs * regs, long error_code)
{
if (regs->eflags & VM_MASK)
goto gp_in_vm86;
/*
虚拟8086下发生的异常否
*/
if (!(regs->xcs & 3))
goto gp_in_kernel;
/*
内核发生的异常否
*/
current->tss.error_code = error_code;
current->tss.trap_no = 13;
force_sig(SIGSEGV, current);
/*
用户程序发生的异常，产生异常信号，
根据异常信号的句柄回调处理函数
*/
return;
gp_in_vm86:
lock_kernel();
handle_vm86_fault((struct kernel_vm86_regs *) regs, error_code);
/*
虚拟8086的处理
*/
unlock_kernel();
return;
gp_in_kernel:
{
unsigned long fixup;
fixup = search_exception_table(regs->eip);
/*
根据异常时的eip搜索异常结构链
找到处理程序地址
*/
if (fixup) {
regs->eip = fixup;
/*
找到异常处理地址，修改中断返回地址，中断返回时跳到异常处理程序处
*/
return;
}
die("general protection fault", regs, error_code);
/*
没找到异常处理程序地址，显示内核异常信息后死机
*/
}
}
搜索异常处理程序代码文件extable.c:
extern const struct exception_table_entry __start___ex_table[];
extern const struct exception_table_entry __stop___ex_table[];
unsigned long search_exception_table(unsigned long addr)
{
unsigned long ret;
#ifndef CONFIG_MODULES
/* There is only the kernel to search. */
ret = search_one_table(__start___ex_table, __stop___ex_table-1, addr);
if (ret) return ret;
#else
/* The kernel is the last "module" -- no need to treat it special. */
struct module *mp;
for (mp = module_list; mp != NULL; mp = mp->next) {
if (mp->ex_table_start == NULL)
continue;
ret = search_one_table(mp->ex_table_start,
mp->ex_table_end - 1, addr);
if (ret) return ret;
}
#endif
return 0;
}
static inline unsigned long
search_one_table(const struct exception_table_entry *first,
const struct exception_table_entry *last,
unsigned long value)
{
while (first <= last) {
const struct exception_table_entry *mid;
long diff;
mid = (last - first) / 2 + first;
diff = mid->insn - value;
if (diff == 0)
return mid->fixup;
else if (diff <0)
first = mid+1;
else
last = mid-1;
}
return 0;
}
看看上面搜索异常处理程序的算法就知道了，有个异常模块表，保存会发生异常时候的eip和异常处理程序指针，发生异常的时候就根据异常时候的eip搜索表里面的eip，发现相等就找到了异常处理指针。这是什么意思呢，就是说你编写的内核程序必须精确的知道哪条指令可能会发生异常，要求真够高的。想想windows下的异常编程是多么轻松？程序员只需要知道哪一段程序可能出现异常，就只需要一个括号一个异常语句保护这段程序就是了。
光看上面算法可能对其这异常的处理还没怎么有感性认识，那么我们再看看其内核的异常形式、编写方式吧。
下面我们再看看exception.txt的一些说明:
查看内核模块表：
> Sections:
> Idx Name Size VMA LMA File off Algn
> 0 .text 00098f40 c0100000 c0100000 00001000 2**4
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 1 .fixup 000016bc c0198f40 c0198f40 00099f40 2**0
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 2 .rodata 0000f127 c019a5fc c019a5fc 0009b5fc 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 3 __ex_table 000015c0 c01a9724 c01a9724 000aa724 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 4 .data 0000ea58 c01abcf0 c01abcf0 000abcf0 2**4
> CONTENTS, ALLOC, LOAD, DATA
> 5 .bss 00018e21 c01ba748 c01ba748 000ba748 2**2
> ALLOC
> 6 .comment 00000ec4 00000000 00000000 000ba748 2**0
> CONTENTS, READONLY
> 7 .note 00001068 00000ec4 00000ec4 000bb60c 2**0
> CONTENTS, READONLY
看模块__ex_table就是会出现异常的一些程序eip，fixup就是相应的异常处理程序地址。这显然这异常结构是静态的，与windows动态的链表形式有很大的分别。
再看看异常程序的编写，下面是get_user(c, buf)的一段代码：
switch ((sizeof(*(buf)))) {
case 1:
__asm__ __volatile__(
"1: mov" "b" " %2,%" "b" "1n"
/*
这语句可能发生异常
*/
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
/*
异常处理程序
*/
" xor" "b" " %" "b" "1,%" "b" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n"
" .long 1b,3bn"
/*
1b,3b就是对应1：，3：就是可能会发生异常的eip与异常后的处理程序
*/
".text" : "=r"(__gu_err), "=q" (__gu_val): "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err )) ;
break;
看看上面代码，是不是要严格的知道哪条指令可能会产生异常？你也清楚了linux整个异常处理方式了吧。编写是不是也太麻烦？弄不好你不知道哪条语句会发生异常，而执行到那发生了异常，那么系统就会出现可怕的异常提示后死机了。
这点完全可以用一个链表的方式处理，哪段程序可能发生异常，就安装异常处理程序指针到链表里面，执行完这段代码就从链表里面删除这段异常处理代码指针，如果这段代码发生异常，系统的异常中断处理程序只需要调用这链表的异常处理程序指针就是了，不要匹配发生异常时的eip。linux的处理主要是因为处理成了一个静态的链表，而不是动态的，这就使得需要确定发生异常时的异常处理指针，所以就增加了一个检测eip完全相等的匹配条件，而这就造成了程序编写上的苦难。这处理方式不用我说你就会知道是多么的糟糕了，弄不好就会留下一些会让系统崩溃的地方呢。
附exception.txt:
Kernel level exception handling in Linux 2.1.8
Commentary by Joerg Pommnitz
When a process runs in kernel mode, it often has to access user
mode memory whose address has been passed by an untrusted program.
To protect itself the kernel has to verify this address.
In older versions of Linux this was done with the
int verify_area(int type, const void * addr, unsigned long size)
function.
This function verified that the memory area starting at address
addr and of size size was accessible for the operation specified
in type (read or write). To do this, verify_read had to look up the
virtual memory area (vma) that contained the address addr. In the
normal case (correctly working program), this test was successful.
It only failed for a few buggy programs. In some kernel profiling
tests, this normally unneeded verification used up a considerable
amount of time.
To overcome this situation, Linus decided to let the virtual memory
hardware present in every Linux-capable CPU handle this test.
How does this work?
Whenever the kernel tries to access an address that is currently not
accessible, the CPU generates a page fault exception and calls the
page fault handler
void do_page_fault(struct pt_regs *regs, unsigned long error_code)
in arch/i386/mm/fault.c. The parameters on the stack are set up by
the low level assembly glue in arch/i386/kernel/entry.S. The parameter
regs is a pointer to the saved registers on the stack, error_code
contains a reason code for the exception.
do_page_fault first obtains the unaccessible address from the CPU
control register CR2. If the address is within the virtual address
space of the process, the fault probably occurred, because the page
was not swapped in, write protected or something similar. However,
we are interested in the other case: the address is not valid, there
is no vma that contains this address. In this case, the kernel jumps
to the bad_area label.
There it uses the address of the instruction that caused the exception
(i.e. regs->eip) to find an address where the execution can continue
(fixup). If this search is successful, the fault handler modifies the
return address (again regs->eip) and returns. The execution will
continue at the address in fixup.
Where does fixup point to?
Since we jump to the contents of fixup, fixup obviously points
to executable code. This code is hidden inside the user access macros.
I have picked the get_user macro defined in include/asm/uaccess.h as an
example. The definition is somewhat hard to follow, so let's peek at
the code generated by the preprocessor and the compiler. I selected
the get_user call in drivers/char/console.c for a detailed examination.
The original code in console.c line 1405:
get_user(c, buf);
The preprocessor output (edited to become somewhat readable):
(
{
long __gu_err = - 14 , __gu_val = 0;
const __typeof__(*( ( buf ) )) *__gu_addr = ((buf));
if (((((0 + current_set[0])->tss.segment) == 0x18 ) ||
(((sizeof(*(buf))) <= 0xC0000000UL) &&
((unsigned long)(__gu_addr ) <= 0xC0000000UL - (sizeof(*(buf)))))))
do {
__gu_err = 0;
switch ((sizeof(*(buf)))) {
case 1:
__asm__ __volatile__(
"1: mov" "b" " %2,%" "b" "1n"
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
" xor" "b" " %" "b" "1,%" "b" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n"
" .long 1b,3bn"
".text" : "=r"(__gu_err), "=q" (__gu_val): "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err )) ;
break;
case 2:
__asm__ __volatile__(
"1: mov" "w" " %2,%" "w" "1n"
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
" xor" "w" " %" "w" "1,%" "w" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n"
" .long 1b,3bn"
".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"( __gu_err ));
break;
case 4:
__asm__ __volatile__(
"1: mov" "l" " %2,%" "" "1n"
"2:n"
".section .fixup,"ax"n"
"3: movl %3,%0n"
" xor" "l" " %" "" "1,%" "" "1n"
" jmp 2bn"
".section __ex_table,"a"n"
" .align 4n" " .long 1b,3bn"
".text" : "=r"(__gu_err), "=r" (__gu_val) : "m"((*(struct __large_struct *)
( __gu_addr )) ), "i"(- 14 ), "0"(__gu_err));
break;
default:
(__gu_val) = __get_user_bad();
}
} while (0) ;
((c)) = (__typeof__(*((buf))))__gu_val;
__gu_err;
}
);
WOW! Black GCC/assembly magic. This is impossible to follow, so let's
see what code gcc generates:
> xorl %edx,%edx
> movl current_set,%eax
> cmpl $24,788(%eax)
> je .L1424
> cmpl $-1073741825,64(%esp)
> ja .L1423
> .L1424:
> movl %edx,%eax
> movl 64(%esp),%ebx
> #APP
> 1: movb (%ebx),%dl /* this is the actual user access */
> 2:
> .section .fixup,"ax"
> 3: movl $-14,%eax
> xorb %dl,%dl
> jmp 2b
> .section __ex_table,"a"
> .align 4
> .long 1b,3b
> .text
> #NO_APP
> .L1423:
> movzbl %dl,%esi
The optimizer does a good job and gives us something we can actually
understand. Can we? The actual user access is quite obvious. Thanks
to the unified address space we can just access the address in user
memory. But what does the .section stuff do?????
To understand this we have to look at the final kernel:
> objdump --section-headers vmlinux
>
> vmlinux: file format elf32-i386
>
> Sections:
> Idx Name Size VMA LMA File off Algn
> 0 .text 00098f40 c0100000 c0100000 00001000 2**4
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 1 .fixup 000016bc c0198f40 c0198f40 00099f40 2**0
> CONTENTS, ALLOC, LOAD, READONLY, CODE
> 2 .rodata 0000f127 c019a5fc c019a5fc 0009b5fc 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 3 __ex_table 000015c0 c01a9724 c01a9724 000aa724 2**2
> CONTENTS, ALLOC, LOAD, READONLY, DATA
> 4 .data 0000ea58 c01abcf0 c01abcf0 000abcf0 2**4
> CONTENTS, ALLOC, LOAD, DATA
> 5 .bss 00018e21 c01ba748 c01ba748 000ba748 2**2
> ALLOC
> 6 .comment 00000ec4 00000000 00000000 000ba748 2**0
> CONTENTS, READONLY
> 7 .note 00001068 00000ec4 00000ec4 000bb60c 2**0
> CONTENTS, READONLY
There are obviously 2 non standard ELF sections in the generated object
file. But first we want to find out what happened to our code in the
final kernel executable:
> objdump --disassemble --section=.text vmlinux
>
> c017e785 xorl %edx,%edx
> c017e787 movl 0xc01c7bec,%eax
> c017e78c cmpl $0x18,0x314(%eax)
> c017e793 je c017e79f
> c017e795 cmpl $0xbfffffff,0x40(%esp,1)
> c017e79d ja c017e7a7
> c017e79f movl %edx,%eax
> c017e7a1 movl 0x40(%esp,1),%ebx
> c017e7a5 movb (%ebx),%dl
> c017e7a7 movzbl %dl,%esi
The whole user memory access is reduced to 10 x86 machine instructions.
The instructions bracketed in the .section directives are not longer
in the normal execution path. They are located in a different section
of the executable file:
> objdump --disassemble --section=.fixup vmlinux
>
> c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
> c0199ffa <.fixup+10ba> xorb %dl,%dl
> c0199ffc <.fixup+10bc> jmp c017e7a7
And finally:
> objdump --full-contents --section=__ex_table vmlinux
>
> c01aa7c4 93c017c0 e09f19c0 97c017c0 99c017c0 ................
> c01aa7d4 f6c217c0 e99f19c0 a5e717c0 f59f19c0 ................
> c01aa7e4 080a18c0 01a019c0 0a0a18c0 04a019c0 ................
or in human readable byte order:
> c01aa7c4 c017c093 c0199fe0 c017c097 c017c099 ................
> c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
^^^^^^^^^^^^^^^^^
this is the interesting part!
> c01aa7e4 c0180a08 c019a001 c0180a0a c019a004 ................
What happened? The assembly directives
.section .fixup,"ax"
.section __ex_table,"a"
told the assembler to move the following code to the specified
sections in the ELF object file. So the instructions
3: movl $-14,%eax
xorb %dl,%dl
jmp 2b
ended up in the .fixup section of the object file and the addresses
.long 1b,3b
ended up in the __ex_table section of the object file. 1b and 3b
are local labels. The local label 1b (1b stands for next label 1
backward) is the address of the instruction that might fault, i.e.
in our case the address of the label 1 is c017e7a5:
the original assembly code: > 1: movb (%ebx),%dl
and linked in vmlinux : > c017e7a5 movb (%ebx),%dl
The local label 3 (backwards again) is the address of the code to handle
the fault, in our case the actual value is c0199ff5:
the original assembly code: > 3: movl $-14,%eax
and linked in vmlinux : > c0199ff5 <.fixup+10b5> movl $0xfffffff2,%eax
The assembly code
> .section __ex_table,"a"
> .align 4
> .long 1b,3b
becomes the value pair
> c01aa7d4 c017c2f6 c0199fe9 c017e7a5 c0199ff5 ................
^this is ^this is
1b 3b
c017e7a5,c0199ff5 in the exception table of the kernel.
So, what actually happens if a fault from kernel mode with no suitable
vma occurs?
1.) access to invalid address:
> c017e7a5 movb (%ebx),%dl
2.) MMU generates exception
3.) CPU calls do_page_fault
4.) do page fault calls search_exception_table (regs->eip == c017e7a5);
5.) search_exception_table looks up the address c017e7a5 in the
exception table (i.e. the contents of the ELF section __ex_table)
and returns the address of the associated fault handle code c0199ff5.
6.) do_page_fault modifies its own return address to point to the fault
handle code and returns.
7.) execution continues in the fault handling code.
8.) 8a) EAX becomes -EFAULT (== -14)
8b) DL becomes zero (the value we "read" from user space)
8c) execution continues at local label 2 (address of the
instruction immediately after the faulting user access).
The steps 8a to 8c in a certain way emulate the faulting instruction.
That's it, mostly. If you look at our example, you might ask why
we set EAX to -EFAULT in the exception handler code. Well, the
get_user macro actually returns a value: 0, if the user access was
successful, -EFAULT on failure. Our original code did not test this
return value, however the inline assembly code in get_user tries to
return -EFAULT. GCC selected EAX to return this value.
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