How noexecstack became a Stack of Confusion

intro

Sometimes, what we perceive as a constant in our programming environments can undergo unexpected shifts, challenging our assumptions. In my journey to understand the mechanics behind stack overflow exploits, I encountered such a shift when grappling with the intricacies of the stack. Initially, as I delved into these techniques using machines devoid of MMUs, namely, plain m68k and x86 real mode, I paid little heed to memory flags. In those days, hackers could seamlessly inject binary payloads onto the stack, redirect program flow to the designated stack address housing their payloads, and execute their exploits with ease.

However, after setting aside these experiments for a time and revisiting them on early Linux machines, I encountered a surprising obstacle around 2005: the once-reliable technique suddenly ceased to function. Upon investigation, I came to the realization that assuming the executability of the stack was no longer tenable. Henceforth, I found myself grappling with the repercussions of this change, as the default behavior of compilers had shifted to render the stack non-executable. Or so I believed, until a recent inquiry from a client prompted me to revisit this assumption, revealing a truth starkly different from my prior expectations.

chapter 1 - What it seems like

So, what do we have here? Since 2005, something peculiar has emerged. When compiling a simple, trivial program using the C compiler, we observe the following:

$ echo -e "#include <stdio.h>\nint main(){printf(\"hello\\\n\");}"| gcc -x c -o hello - ;readelf -l hello Elf file type is EXEC (Executable file) Entry point 0x4004a0 There are 9 program headers, starting at offset 64 Program Headers: Type Offset VirtAddr PhysAddr FileSiz MemSiz Flags Align PHDR 0x0000000000000040 0x0000000000400040 0x0000000000400040 0x00000000000001f8 0x00000000000001f8 R 0x8 INTERP 0x0000000000000238 0x0000000000400238 0x0000000000400238 0x000000000000001c 0x000000000000001c R 0x1 [Requesting program interpreter: /lib64/ld-linux-x86-64.so.2] LOAD 0x0000000000000000 0x0000000000400000 0x0000000000400000 0x0000000000000768 0x0000000000000768 R E 0x200000 LOAD 0x0000000000000e00 0x0000000000600e00 0x0000000000600e00 0x0000000000000224 0x0000000000000228 RW 0x200000 DYNAMIC 0x0000000000000e10 0x0000000000600e10 0x0000000000600e10 0x00000000000001d0 0x00000000000001d0 RW 0x8 NOTE 0x0000000000000254 0x0000000000400254 0x0000000000400254 0x0000000000000044 0x0000000000000044 R 0x4 GNU_EH_FRAME 0x0000000000000640 0x0000000000400640 0x0000000000400640 0x000000000000003c 0x000000000000003c R 0x4 GNU_STACK 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 RW 0x10 GNU_RELRO 0x0000000000000e00 0x0000000000600e00 0x0000000000600e00 0x0000000000000200 0x0000000000000200 R 0x1 Section to Segment mapping: Segment Sections... 00 01 .interp 02 .interp .note.ABI-tag .note.gnu.build-id .gnu.hash .dynsym .dynstr .gnu.version .gnu.version_r .rela.dyn .rela.plt .init .plt .text .fini .rodata .eh_frame_hdr .eh_frame 03 .init_array .fini_array .dynamic .got .got.plt .data .bss 04 .dynamic 05 .note.ABI-tag .note.gnu.build-id 06 .eh_frame_hdr 07 08 .init_array .fini_array .dynamic .got

Not much needs to be said; the stack lacks an executable flag: RW in the GNU_STACK section. Any attempt to execute code from this space inevitably results in a graceful crash, marked by the familiar segmentation fault (SIGSEGV).

Conversely, if our intention is to create an executable stack, we must explicitly instruct the compiler to do so.

$ echo -e "#include <stdio.h>\nint main(){printf(\"hello\\\n\");}"| gcc -x c -z execstack -o hello - ;readelf -l hello Elf file type is EXEC (Executable file) Entry point 0x4004a0 There are 9 program headers, starting at offset 64 Program Headers: Type Offset VirtAddr PhysAddr FileSiz MemSiz Flags Align PHDR 0x0000000000000040 0x0000000000400040 0x0000000000400040 0x00000000000001f8 0x00000000000001f8 R 0x8 INTERP 0x0000000000000238 0x0000000000400238 0x0000000000400238 0x000000000000001c 0x000000000000001c R 0x1 [Requesting program interpreter: /lib64/ld-linux-x86-64.so.2] LOAD 0x0000000000000000 0x0000000000400000 0x0000000000400000 0x0000000000000768 0x0000000000000768 R E 0x200000 LOAD 0x0000000000000e00 0x0000000000600e00 0x0000000000600e00 0x0000000000000224 0x0000000000000228 RW 0x200000 DYNAMIC 0x0000000000000e10 0x0000000000600e10 0x0000000000600e10 0x00000000000001d0 0x00000000000001d0 RW 0x8 NOTE 0x0000000000000254 0x0000000000400254 0x0000000000400254 0x0000000000000044 0x0000000000000044 R 0x4 GNU_EH_FRAME 0x0000000000000640 0x0000000000400640 0x0000000000400640 0x000000000000003c 0x000000000000003c R 0x4 GNU_STACK 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 RWE 0x10 GNU_RELRO 0x0000000000000e00 0x0000000000600e00 0x0000000000600e00 0x0000000000000200 0x0000000000000200 R 0x1 Section to Segment mapping: Segment Sections... 00 01 .interp 02 .interp .note.ABI-tag .note.gnu.build-id .gnu.hash .dynsym .dynstr .gnu.version .gnu.version_r .rela.dyn .rela.plt .init .plt .text .fini .rodata .eh_frame_hdr .eh_frame 03 .init_array .fini_array .dynamic .got .got.plt .data .bss 04 .dynamic 05 .note.ABI-tag .note.gnu.build-id 06 .eh_frame_hdr 07 08 .init_array .fini_array .dynamic .got

Upon inspection, RWE in the GNU_STACK section, we confirm the presence of an executable stack.

In summary, the probability of encountering a new binary with an executable stack in contemporary settings is close to zero. Such instances may occur only if someone utilizes an outdated compiler or requires an executable stack for specific reasons. So, why would anyone desire an executable stack?

Perhaps solely to revisit the methods employed in old-fashioned stack overflow exploits!

Chapter 2 - Things are never a easy as they seems

Recently, a customer posed what initially appeared to be a trivial question: “What flag should I use to ensure that the stack remains non-executable?” I brushed it off as a simple matter, assuming that no action was needed since it was the default behavior.

However, the response from a knowledgeable individual surprised me: simply use -z nostackexec. This prompted me to question why such a flag even existed. After all, if the default behavior is to have a non-executable stack, what purpose does this flag serve?

Reflecting on past encounters with this flag, I had rationalized its existence by speculating, “Perhaps it’s necessary for exotic architectures where the default is to have an executable stack”.

However, I soon realized that the reality is far more complex than it initially seemed.

Let’s begin by clarifying: compilers do not manipulate stack flags; this task falls under the responsibility of the linker. The final executable is created by linking together all the object files generated by the compiler.

During the creation of ELF sections, the linker scans the input files for a specific section named .note.GNU-stack. This section conveys whether an executable stack is required or not.

According to the linker’s manual page, if an input file lacks a .note.GNU-stack section, then the default behavior is architecture-specific.

As I couldn’t find where this default behavior is specified, let’s conduct a couple of tests. You can find a collection of tests I’ve prepared in this repository.

Consider the gcc/asm_function executable file, which is a simple C executable that includes a basic function from an assembly file. Below is the relevant portion of the Makefile used to build it:

gcc/asm_function.o: src/asm_function.S gcc -g -c -o gcc/asm_function.o src/asm_function.S gcc/test_asm.o: src/test_asm.c gcc -g -c -o gcc/test_asm.o src/test_asm.c gcc/test_asm: gcc/test_asm.o gcc/asm_function.o gcc -g gcc/test_asm.o gcc/asm_function.o -o gcc/asm_function

Upon examining the generated object file, you’ll notice the absence of the .note.GNU-stack section. However, upon inspecting the resultant executable, you’ll observe that the stack is indeed marked as executable.

$ readelf -S gcc/asm_function.o There are 15 section headers, starting at offset 0x3b8: Section Headers: [Nr] Name Type Address Offset Size EntSize Flags Link Info Align [ 0] NULL 0000000000000000 00000000 0000000000000000 0000000000000000 0 0 0 [ 1] .text PROGBITS 0000000000000000 00000040 0000000000000006 0000000000000000 AX 0 0 1 [ 2] .data PROGBITS 0000000000000000 00000046 0000000000000000 0000000000000000 WA 0 0 1 [ 3] .bss NOBITS 0000000000000000 00000046 0000000000000000 0000000000000000 WA 0 0 1 [ 4] .debug_line PROGBITS 0000000000000000 00000046 0000000000000045 0000000000000000 0 0 1 [ 5] .rela.debug_line RELA 0000000000000000 00000248 0000000000000018 0000000000000018 I 12 4 8 [ 6] .debug_info PROGBITS 0000000000000000 0000008b 000000000000002e 0000000000000000 0 0 1 [ 7] .rela.debug_info RELA 0000000000000000 00000260 00000000000000a8 0000000000000018 I 12 6 8 [ 8] .debug_abbrev PROGBITS 0000000000000000 000000b9 0000000000000014 0000000000000000 0 0 1 [ 9] .debug_aranges PROGBITS 0000000000000000 000000d0 0000000000000030 0000000000000000 0 0 16 [10] .rela.debug_arang RELA 0000000000000000 00000308 0000000000000030 0000000000000018 I 12 9 8 [11] .debug_str PROGBITS 0000000000000000 00000100 0000000000000045 0000000000000001 MS 0 0 1 [12] .symtab SYMTAB 0000000000000000 00000148 00000000000000f0 0000000000000018 13 9 8 [13] .strtab STRTAB 0000000000000000 00000238 0000000000000009 0000000000000000 0 0 1 [14] .shstrtab STRTAB 0000000000000000 00000338 000000000000007b 0000000000000000 0 0 1 Key to Flags: W (write), A (alloc), X (execute), M (merge), S (strings), I (info), L (link order), O (extra OS processing required), G (group), T (TLS), C (compressed), x (unknown), o (OS specific), E (exclude), l (large), p (processor specific) $ readelf -l gcc/asm_function Elf file type is DYN (Shared object file) Entry point 0x1040 There are 11 program headers, starting at offset 64 Program Headers: Type Offset VirtAddr PhysAddr FileSiz MemSiz Flags Align PHDR 0x0000000000000040 0x0000000000000040 0x0000000000000040 0x0000000000000268 0x0000000000000268 R 0x8 INTERP 0x00000000000002a8 0x00000000000002a8 0x00000000000002a8 0x000000000000001c 0x000000000000001c R 0x1 [Requesting program interpreter: /lib64/ld-linux-x86-64.so.2] LOAD 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000530 0x0000000000000530 R 0x1000 LOAD 0x0000000000001000 0x0000000000001000 0x0000000000001000 0x00000000000001d5 0x00000000000001d5 R E 0x1000 LOAD 0x0000000000002000 0x0000000000002000 0x0000000000002000 0x0000000000000130 0x0000000000000130 R 0x1000 LOAD 0x0000000000002df0 0x0000000000003df0 0x0000000000003df0 0x0000000000000220 0x0000000000000228 RW 0x1000 DYNAMIC 0x0000000000002e00 0x0000000000003e00 0x0000000000003e00 0x00000000000001c0 0x00000000000001c0 RW 0x8 NOTE 0x00000000000002c4 0x00000000000002c4 0x00000000000002c4 0x0000000000000044 0x0000000000000044 R 0x4 GNU_EH_FRAME 0x0000000000002004 0x0000000000002004 0x0000000000002004 0x000000000000003c 0x000000000000003c R 0x4 GNU_STACK 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 RWE 0x10 GNU_RELRO 0x0000000000002df0 0x0000000000003df0 0x0000000000003df0 0x0000000000000210 0x0000000000000210 R 0x1 Section to Segment mapping: Segment Sections... 00 01 .interp 02 .interp .note.gnu.build-id .note.ABI-tag .gnu.hash .dynsym .dynstr .gnu.version .gnu.version_r .rela.dyn 03 .init .plt .plt.got .text .fini 04 .rodata .eh_frame_hdr .eh_frame 05 .init_array .fini_array .dynamic .got .data .bss 06 .dynamic 07 .note.gnu.build-id .note.ABI-tag 08 .eh_frame_hdr 09 10 .init_array .fini_array .dynamic .got

This suggests that the default for x86_64 architecture is executable stack. Doing the same for aarch64, produces the followings:

$ readelf -S gcc/asm_function.aarch64.o There are 15 section headers, starting at offset 0x400: Section Headers: [Nr] Name Type Address Offset Size EntSize Flags Link Info Align [ 0] NULL 0000000000000000 00000000 0000000000000000 0000000000000000 0 0 0 [ 1] .text PROGBITS 0000000000000000 00000040 0000000000000010 0000000000000000 AX 0 0 8 [ 2] .data PROGBITS 0000000000000000 00000050 0000000000000000 0000000000000000 WA 0 0 1 [ 3] .bss NOBITS 0000000000000000 00000050 0000000000000000 0000000000000000 WA 0 0 1 [ 4] .debug_line PROGBITS 0000000000000000 00000050 000000000000004c 0000000000000000 0 0 1 [ 5] .rela.debug_line RELA 0000000000000000 00000290 0000000000000018 0000000000000018 I 12 4 8 [ 6] .debug_info PROGBITS 0000000000000000 0000009c 000000000000002e 0000000000000000 0 0 1 [ 7] .rela.debug_info RELA 0000000000000000 000002a8 00000000000000a8 0000000000000018 I 12 6 8 [ 8] .debug_abbrev PROGBITS 0000000000000000 000000ca 0000000000000014 0000000000000000 0 0 1 [ 9] .debug_aranges PROGBITS 0000000000000000 000000e0 0000000000000030 0000000000000000 0 0 16 [10] .rela.debug_arang RELA 0000000000000000 00000350 0000000000000030 0000000000000018 I 12 9 8 [11] .debug_str PROGBITS 0000000000000000 00000110 000000000000004d 0000000000000001 MS 0 0 1 [12] .symtab SYMTAB 0000000000000000 00000160 0000000000000120 0000000000000018 13 11 8 [13] .strtab STRTAB 0000000000000000 00000280 000000000000000f 0000000000000000 0 0 1 [14] .shstrtab STRTAB 0000000000000000 00000380 000000000000007b 0000000000000000 0 0 1 Key to Flags: W (write), A (alloc), X (execute), M (merge), S (strings), I (info), L (link order), O (extra OS processing required), G (group), T (TLS), C (compressed), x (unknown), o (OS specific), E (exclude), p (processor specific) $ readelf -l gcc/asm_function.aarch64 Elf file type is DYN (Shared object file) Entry point 0x610 There are 9 program headers, starting at offset 64 Program Headers: Type Offset VirtAddr PhysAddr FileSiz MemSiz Flags Align PHDR 0x0000000000000040 0x0000000000000040 0x0000000000000040 0x00000000000001f8 0x00000000000001f8 R 0x8 INTERP 0x0000000000000238 0x0000000000000238 0x0000000000000238 0x000000000000001b 0x000000000000001b R 0x1 [Requesting program interpreter: /lib/ld-linux-aarch64.so.1] LOAD 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x000000000000090c 0x000000000000090c R E 0x10000 LOAD 0x0000000000000d88 0x0000000000010d88 0x0000000000010d88 0x0000000000000288 0x0000000000000290 RW 0x10000 DYNAMIC 0x0000000000000d98 0x0000000000010d98 0x0000000000010d98 0x00000000000001f0 0x00000000000001f0 RW 0x8 NOTE 0x0000000000000254 0x0000000000000254 0x0000000000000254 0x0000000000000044 0x0000000000000044 R 0x4 GNU_EH_FRAME 0x00000000000007e0 0x00000000000007e0 0x00000000000007e0 0x0000000000000044 0x0000000000000044 R 0x4 GNU_STACK 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 RW 0x10 GNU_RELRO 0x0000000000000d88 0x0000000000010d88 0x0000000000010d88 0x0000000000000278 0x0000000000000278 R 0x1 Section to Segment mapping: Segment Sections... 00 01 .interp 02 .interp .note.gnu.build-id .note.ABI-tag .gnu.hash .dynsym .dynstr .gnu.version .gnu.version_r .rela.dyn .rela.plt .init .plt .text .fini .rodata .eh_frame_hdr .eh_frame 03 .init_array .fini_array .dynamic .got .data .bss 04 .dynamic 05 .note.gnu.build-id .note.ABI-tag 06 .eh_frame_hdr 07 08 .init_array .fini_array .dynamic .got

Which is the opposite, non-executable stack, suggesting that the default for this architecture is to have the stack not executable.

chapter 3 - is it all?

Returning to the original topic, based on the previous chapter observations, it is possible to deduce that the two architectures, x86_64 and aarch64, have different defaults regarding executable stacks. But is this the extent of the matter?

Apparently not. There are instances where the compiler needs to generate code and execute it, often utilizing the stack for this purpose. In such cases, the resulting executable file will indeed have an executable stack.

There might be other cases out there, but after a thorough search, I couldn’t find anything except for the GCC GNU extension “nested functions.” It’s possible that not many people are aware of this feature - I certainly wasn’t until recently. However, it appears that nested functions can be implemented in C, but only when using GCC, clang does not support them.

Nested functions are functions defined within the body of another function. These inner functions have access to the variables and parameters of the enclosing function and can only be invoked within its scope. GCC allows them to exist, but for them to work, the stack needs to be executable, at least when they are called indirectly from another function.

Let’s consider an example:

int nested_carrier(int a, int b, int n) { int loc_var = n; int multiply2(int z) { return z + z + loc_var; } return sum_func(multiply2, a, b); }

In this function, multiply2 is passed to be executed by the external function sum_func. Now, let’s examine the assembly implementation of nested_carrier.

┌ 151: dbg.nested_carrier (int64_t arg1, int64_t arg2, int64_t arg3, int64_t arg_10h); │ ; arg int64_t arg1 @ rdi │ ; arg int64_t arg2 @ rsi │ ; arg int64_t arg3 @ rdx │ ; arg int64_t arg_10h @ rbp+0x10 │ ; var int z @ rbp-0x4 │ ; var int64_t canary @ rbp-0x8 │ ; var int64_t var_10h @ rbp-0x10 │ ; var int loc_var @ rbp-0x30 │ ; var int a @ rbp-0x34 │ ; var int b @ rbp-0x38 │ ; var int n @ rbp-0x3c │ 0x00001187 f30f1efa endbr64 ; nested_local.c:5 int nested_carrier (int a, int b, int n) { │ ; int nested_carrier(int a,int b,int n); │ 0x0000118b 55 push rbp │ 0x0000118c 4889e5 mov rbp, rsp │ 0x0000118f 4883ec40 sub rsp, 0x40 │ 0x00001193 897dcc mov dword [a], edi ; arg1 │ 0x00001196 8975c8 mov dword [b], esi ; arg2 │ 0x00001199 8955c4 mov dword [n], edx ; arg3 │ 0x0000119c 64488b0425.. mov rax, qword fs:[0x28] │ 0x000011a5 488945f8 mov qword [canary], rax ; Just bought my self a new canary │ 0x000011a9 31c0 xor eax, eax │ 0x000011ab 488d4510 lea rax, [arg_10h] │ 0x000011af 488945f0 mov qword [var_10h], rax │ 0x000011b3 488d45d0 lea rax, [loc_var] │ 0x000011b7 4883c004 add rax, 4 │ 0x000011bb 488d55d0 lea rdx, [loc_var] │ 0x000011bf c700f30f1efa mov dword [rax], 0xfa1e0ff3 ; Here it is writing the trampoline, note the endbr64 opcode │ 0x000011c5 66c7400449bb mov word [rax + 4], 0xbb49 ; it stores in the stack │ 0x000011cb 488d0d97ff.. lea rcx, [dbg.multiply2] ; as the multiply2 address │ 0x000011d2 48894806 mov qword [rax + 6], rcx │ 0x000011d6 66c7400e49ba mov word [rax + 0xe], 0xba49 ; another opcode │ 0x000011dc 48895010 mov qword [rax + 0x10], rdx ; this is the base address to locate parent local vars │ 0x000011e0 c7401849ff.. mov dword [rax + 0x18], 0x90e3ff49 ; more opcodes │ 0x000011e7 8b45c4 mov eax, dword [n] ; nested_local.c:6 int loc_var = n; │ 0x000011ea 8945d0 mov dword [loc_var], eax │ 0x000011ed 488d45d0 lea rax, [loc_var] ; nested_local.c:8 return sum_func (multiply2, a, b); │ 0x000011f1 4883c004 add rax, 4 │ 0x000011f5 4889c1 mov rcx, rax ; save trampoline address │ 0x000011f8 8b55c8 mov edx, dword [b] ; int64_t arg3 = b │ 0x000011fb 8b45cc mov eax, dword [a] │ 0x000011fe 89c6 mov esi, eax ; int64_t arg2 = a │ 0x00001200 4889cf mov rdi, rcx ; int64_t arg1 = trampoline address! │ 0x00001203 e847000000 call dbg.sum_func │ 0x00001208 488b75f8 mov rsi, qword [canary] ; Hey canary, are you there?! │ 0x0000120c 6448333425.. xor rsi, qword fs:[0x28] ; are still you!? │ ┌─< 0x00001215 7405 je 0x121c ; stack overflow check │ │ 0x00001217 e844feffff call sym.imp.__stack_chk_fail ; crash if canary is failing │ └─> 0x0000121c c9 leave └ 0x0000121d c3 ret

Examining this code, we notice some “alien code” added by our trusty compiler friend. Let’s set aside the stack check with canary for now; our current focus is on the trampoline it’s constructing to facilitate the external call. Within the function body, we can clearly see the trampoline being constructed, followed by the point at which the trampoline address is utilized for the external function call.

(gdb) x/10i $pc => 0x7fffffffdcc0: endbr64 0x7fffffffdcc4: movabs $0x555555555169,%r11 0x7fffffffdcce: movabs $0x7fffffffdcc0,%r10 0x7fffffffdcd8: rex.WB jmpq *%r11

Let’s delve into how the trampoline is constructed using our buddy GDB. We’ll break it down into four instructions:

1. The endbr64 instruction was introduced as part of the Intel Control-flow Enforcement Technology (CET) extension. Don’t confuse it with Cache Allocation Technology (CAT), another CPU feature. Phew, the acronyms are piling up! Anyway, this instruction isn’t pertinent to our analysis; it’s included because the machine executing this code expects it to be present. The endbr64 instruction marks the end of a code sequence and helps prevent ROP gadgets from being chained together.
2. movabs $0x555555555169,%r11: This instruction loads our target function address, multiply2, into register r11.
3. movabs $0x7fffffffdcc0,%r10: Let’s recall the x86_64 ABI: Parameters to functions are passed in the registers rdi, rsi, rdx, rcx, r8, r9, and additional values are passed on the stack in reverse order. This instruction deviates from the conventional ABI, using a register r10, to pass the base address for the parent’s local variables.
4. rex.WB jmpq *%r11: This is a straightforward indirect call that we know will lead us to address 0x555555555169, corresponding to the multiply2 function.

Now that we are aware of at least one other scenario where the compiler may necessitate an executable stack, let’s explore how this is reflected in the executable:

$ readelf -l gcc/nested_local Elf file type is DYN (Shared object file) Entry point 0x1080 There are 13 program headers, starting at offset 64 Program Headers: Type Offset VirtAddr PhysAddr FileSiz MemSiz Flags Align PHDR 0x0000000000000040 0x0000000000000040 0x0000000000000040 0x00000000000002d8 0x00000000000002d8 R 0x8 INTERP 0x0000000000000318 0x0000000000000318 0x0000000000000318 0x000000000000001c 0x000000000000001c R 0x1 [Requesting program interpreter: /lib64/ld-linux-x86-64.so.2] LOAD 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000658 0x0000000000000658 R 0x1000 LOAD 0x0000000000001000 0x0000000000001000 0x0000000000001000 0x0000000000000315 0x0000000000000315 R E 0x1000 LOAD 0x0000000000002000 0x0000000000002000 0x0000000000002000 0x00000000000001e0 0x00000000000001e0 R 0x1000 LOAD 0x0000000000002db0 0x0000000000003db0 0x0000000000003db0 0x0000000000000260 0x0000000000000268 RW 0x1000 DYNAMIC 0x0000000000002dc0 0x0000000000003dc0 0x0000000000003dc0 0x00000000000001f0 0x00000000000001f0 RW 0x8 NOTE 0x0000000000000338 0x0000000000000338 0x0000000000000338 0x0000000000000020 0x0000000000000020 R 0x8 NOTE 0x0000000000000358 0x0000000000000358 0x0000000000000358 0x0000000000000044 0x0000000000000044 R 0x4 GNU_PROPERTY 0x0000000000000338 0x0000000000000338 0x0000000000000338 0x0000000000000020 0x0000000000000020 R 0x8 GNU_EH_FRAME 0x000000000000201c 0x000000000000201c 0x000000000000201c 0x000000000000005c 0x000000000000005c R 0x4 GNU_STACK 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 RWE 0x10 GNU_RELRO 0x0000000000002db0 0x0000000000003db0 0x0000000000003db0 0x0000000000000250 0x0000000000000250 R 0x1 Section to Segment mapping: Segment Sections... 00 01 .interp 02 .interp .note.gnu.property .note.gnu.build-id .note.ABI-tag .gnu.hash .dynsym .dynstr .gnu.version .gnu.version_r .rela.dyn .rela.plt 03 .init .plt .plt.got .plt.sec .text .fini 04 .rodata .eh_frame_hdr .eh_frame 05 .init_array .fini_array .dynamic .got .data .bss 06 .dynamic 07 .note.gnu.property 08 .note.gnu.build-id .note.ABI-tag 09 .note.gnu.property 10 .eh_frame_hdr 11 12 .init_array .fini_array .dynamic .got $ readelf -l gcc/nested_local.ne Elf file type is DYN (Shared object file) Entry point 0x1080 There are 13 program headers, starting at offset 64 Program Headers: Type Offset VirtAddr PhysAddr FileSiz MemSiz Flags Align PHDR 0x0000000000000040 0x0000000000000040 0x0000000000000040 0x00000000000002d8 0x00000000000002d8 R 0x8 INTERP 0x0000000000000318 0x0000000000000318 0x0000000000000318 0x000000000000001c 0x000000000000001c R 0x1 [Requesting program interpreter: /lib64/ld-linux-x86-64.so.2] LOAD 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000658 0x0000000000000658 R 0x1000 LOAD 0x0000000000001000 0x0000000000001000 0x0000000000001000 0x0000000000000315 0x0000000000000315 R E 0x1000 LOAD 0x0000000000002000 0x0000000000002000 0x0000000000002000 0x00000000000001e0 0x00000000000001e0 R 0x1000 LOAD 0x0000000000002db0 0x0000000000003db0 0x0000000000003db0 0x0000000000000260 0x0000000000000268 RW 0x1000 DYNAMIC 0x0000000000002dc0 0x0000000000003dc0 0x0000000000003dc0 0x00000000000001f0 0x00000000000001f0 RW 0x8 NOTE 0x0000000000000338 0x0000000000000338 0x0000000000000338 0x0000000000000020 0x0000000000000020 R 0x8 NOTE 0x0000000000000358 0x0000000000000358 0x0000000000000358 0x0000000000000044 0x0000000000000044 R 0x4 GNU_PROPERTY 0x0000000000000338 0x0000000000000338 0x0000000000000338 0x0000000000000020 0x0000000000000020 R 0x8 GNU_EH_FRAME 0x000000000000201c 0x000000000000201c 0x000000000000201c 0x000000000000005c 0x000000000000005c R 0x4 GNU_STACK 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 0x0000000000000000 RW 0x10 GNU_RELRO 0x0000000000002db0 0x0000000000003db0 0x0000000000003db0 0x0000000000000250 0x0000000000000250 R 0x1 Section to Segment mapping: Segment Sections... 00 01 .interp 02 .interp .note.gnu.property .note.gnu.build-id .note.ABI-tag .gnu.hash .dynsym .dynstr .gnu.version .gnu.version_r .rela.dyn .rela.plt 03 .init .plt .plt.got .plt.sec .text .fini 04 .rodata .eh_frame_hdr .eh_frame 05 .init_array .fini_array .dynamic .got .data .bss 06 .dynamic 07 .note.gnu.property 08 .note.gnu.build-id .note.ABI-tag 09 .note.gnu.property 10 .eh_frame_hdr 11 12 .init_array .fini_array .dynamic .got

In the following, you can observe the ELF program header table of two executable files, both generated from the same source file, src/nested_local.c, in the repository . They differ because in one instance, I added -z noexecstack to enforce a non-executable stack. This is what’s happen if they get executed:

$ ./gcc/nested_local; echo Fancy calculation (34) alessandro@r5:~/tmp/stack/nested_prt$ ./gcc/nested_local.ne; echo Segmentation fault (core dumped)

Since the trampoline is in the stack, when the second file is executed it crashes because it tries to execute code from the stack. Here’s the proof the crash is caused by it:

$ gdb ./gcc/nested_local.ne GNU gdb (Ubuntu 9.2-0ubuntu1~20.04.1) 9.2 Copyright (C) 2020 Free Software Foundation, Inc. License GPLv3+: GNU GPL version 3 or later <http://gnu.org/licenses/gpl.html> This is free software: you are free to change and redistribute it. There is NO WARRANTY, to the extent permitted by law. Type "show copying" and "show warranty" for details. This GDB was configured as "x86_64-linux-gnu". Type "show configuration" for configuration details. For bug reporting instructions, please see: <http://www.gnu.org/software/gdb/bugs/>. Find the GDB manual and other documentation resources online at: <http://www.gnu.org/software/gdb/documentation/>. For help, type "help". Type "apropos word" to search for commands related to "word"... Reading symbols from ./gcc/nested_local.ne... (gdb) r Starting program: /home/alessandro/tmp/stack/nested_prt/gcc/nested_local.ne Program received signal SIGSEGV, Segmentation fault. 0x00007fffffffdc34 in ?? () (gdb) x/10i $pc => 0x7fffffffdc34: endbr64 0x7fffffffdc38: movabs $0x555555555169,%r11 0x7fffffffdc42: movabs $0x7fffffffdc30,%r10 0x7fffffffdc4c: rex.WB jmpq *%r11 0x7fffffffdc4f: nop 0x7fffffffdc50: jo 0x7fffffffdc2e 0x7fffffffdc52: (bad) 0x7fffffffdc53: (bad) 0x7fffffffdc54: (bad) 0x7fffffffdc55: jg 0x7fffffffdc57 (gdb)

Finally, let’s consider that to further complicate matters, GCC employs different conventions across architectures. Please do not expect this to be straightforward, as it certainly isn’t!

In x86_64, executable ELF files always contain an entry in the program header GNU_STACK, which reflects the actual permissions over the stack. When the linker combines objects to create the executable, it looks at .note.GNU-stack and its contents to set the stack accordingly. If .note.GNU-stack is missing, the stack defaults to executable.

Similarly, in aarch64, executable ELF files always include an entry in the program header GNU_STACK, with flags reflecting the stack’s permissions. The linker examines .note.GNU-stack during the executable creation process to determine the stack’s permissions. If .note.GNU-stack is absent, the stack defaults to non-executable.

On PPC64, executable ELF files only include an entry in the program header GNU_STACK if it needs to be executable; otherwise, it defaults to non-executable.

Conversely, in MIPS32, executable ELF files only have an entry in the program header GNU_STACK if it needs to be non-executable; otherwise, it defaults to executable.

As a final note for this extensive and perhaps tedious discussion on executable stacks, allow me to share what I discovered while verifying this information on a MIPS system.

Look at how some MIPS SoCs do not enforce the stack permissions

epilogue

If you want to make sure your stack is not executable, add -z noexecstack to your compiler's flags.

Linux Syscall Numbers: A Journey through Binary Analysis

Introduction:

Reverse engineering binaries is a crucial activity impacting both the realms of security and safety. When examining a program without the context syscall numbers provide, interpreting the code behavior can be daunting.

Consider this snippet of code devoid of syscall numbers:

0x0000000000000000: 48 C7 C0 01 00 00 00 mov rax, 1 0x0000000000000007: 48 C7 C7 01 00 00 00 mov rdi, 1 0x000000000000000e: 48 C7 C6 2e 00 00 00 mov rsi, 0x2e 0x0000000000000015: 48 C7 C2 0D 00 00 00 mov rdx, 0xd 0x000000000000001c: 0F 05 syscall 0x000000000000001e: 48 C7 C0 3C 00 00 00 mov rax, 0x3c 0x0000000000000025: 48 C7 C7 00 00 00 00 mov rdi, 0 0x000000000000002c: 0F 05 syscall 0x000000000000002e: 48 65 6C insb byte ptr [rdi], dx 0x0000000000000031: 6C insb byte ptr [rdi], dx 0x0000000000000032: 6F outsd dx, dword ptr [rsi] 0x0000000000000033: 20 57 6F and byte ptr [rdi + 0x6f], dl 0x0000000000000036: 72 6C jb 0xa4 0x0000000000000038: 64 21 00 and dword ptr fs:[rax], eax

Without knowledge of syscall numbers, understanding the program’s functionality, particularly its interactions with the operating system, becomes challenging. However, simply providing the context of syscall numbers can dramatically simplify comprehension. What if I tell you that the first syscall is write and the second is exit? Instantainly it became trivial to understand!

0x0000000000000000: 48 C7 C0 01 00 00 00 mov rax, 1 0x0000000000000007: 48 C7 C7 01 00 00 00 mov rdi, 1 0x000000000000000e: 48 C7 C6 2e 00 00 00 mov rsi, 0x2e 0x0000000000000015: 48 C7 C2 0D 00 00 00 mov rdx, 0xd 0x000000000000001c: 0F 05 syscall ; write(1, 0x2e, 13) 0x000000000000001e: 48 C7 C0 3C 00 00 00 mov rax, 0x3c 0x0000000000000025: 48 C7 C7 00 00 00 00 mov rdi, 0 0x000000000000002c: 0F 05 syscall ; exit(0) 0x000000000000002e: 48 65 6C insb byte ptr [rdi], dx ; H e l 0x0000000000000031: 6C insb byte ptr [rdi], dx ; l 0x0000000000000032: 6F outsd dx, dword ptr [rsi] ; o 0x0000000000000033: 20 57 6F and byte ptr [rdi + 0x6f], dl ; \b W o 0x0000000000000036: 72 6C jb 0xa4 ; r l 0x0000000000000038: 64 21 00 and dword ptr fs:[rax], eax ; d ! \0

It become easier to spot that bytes after the exit can’t be code, and that if write references them in a write to 1, stdout, they must be text!

In the snippets I just used as example spot the syscall number is trivial as mov rax, 1 for the first syscall, and mov rax, 0x3c for the second. So the example can just provide an hint of the problem. In a complex code the mov rax in question can be far away from the actual usage and in some convoluted cases, it can be moved from a register to another.

Automating the identification of syscall numbers not only aids reverse engineering for security but also can also enhance code safety. By automating this process, the amount of code that requires manual verification can be restricted. This reduction in manual effort streamlines the verification process, contributing to improved code safety and reliability.

Differences Between Architectures RISC vs CISC:

In the realm of computer architecture, the dynamic between CISC and RISC is akin to a timeless dance. CISC represents the old school, while RISC stands as the new kid on the block — although, truth be told, RISC has been strutting its stuff for at least 30 years now, which is quite a few lifetimes in architectural terms.

RISC architectures boast uniform instruction sizes, treating every instruction with equal importance. This uniformity simplifies the compiler’s task, ensuring impartial treatment of instructions. For instance, in aarch64, a 16-bit literal can effortlessly fit within a single instruction’s operating code, blending seamlessly with others within the 64k limit — a typical boundary for syscall numbers.

Conversely, CISC architectures march to a different beat. Loading an immediate value can entail performance and memory footprint costs. Even with minimal optimizations, varied instructions are often used to achieve the same outcome of loading a register with a specific value.

opcodes	mnemonic
C8 1a 80 D2	MOV X8, #0xd6
01 00 00 D4	SVC 0
RISC: aarch64

opcodes	mnemonic
93 08 e0 05	li a7, 94
0f 05	ecall
RISC: RiscV

opcodes	mnemonic
48 c7 c0 0c 00 00 00	mov rax,0x0C
73 00 00 00	syscall
CISC: x86_64

opcodes	mnemonic
A7 19 01 04	lghi r1, 0x0104
0a 00	svc 0
CISC: s390x

This architectural dichotomy sheds light on the nuances of identifying syscall numbers. RISC architectures, with their uniform instruction structure, facilitate easier identification, while CISC architectures, with their diverse and sometimes costly instructions, present a more formidable challenge.

Guess syscall numbers: How can it be done?

In the early days of computing, CPU designers anticipated the need for system call functionality directly embedded within instruction sets. For instance, the int instruction in 16-bit x86 architectures and traps in m68k systems exemplify this approach. However, as early as the 1980s, operating system developers recognized the limitations of such implementations, as the 256 functions allowed by these instructions proved insufficient. Legacy x86 systems, DOS for instance, differentiated between BIOS and operating system functions using the int number, with the OS typically consolidating its functions under a single int number, such as int 0x21. Today, the impracticality of embedding syscall numbers directly into instruction opcodes is evident, leading newer architectures like RISC-V and x86_64 to eschew this practice.

By hand: How do they do it?

Manual identification of syscall numbers is a cumbersome process, typically undertaken in two steps:

Locating a syscall instruction within the assembly code.

Tracing backward from the syscall instruction to identify registers, particularly the one containing the syscall number.

This meticulous process is essential for determining the specific syscall invoked by the program but highlights the need for automated tools and techniques in modern reverse engineering practices.

Radare ESIL: How it does it?

Radare 2 tackles the challenge of identifying syscall numbers by leveraging its ESIL framework. ESIL serves as a semantic representation of assembly instructions, enabling Radare 2 to evaluate individual instructions and emulate their behavior. Similar to LLVM-IR, ESIL provides a high-level abstraction of CPU opcode semantics, facilitating analysis and interpretation of assembly code.

ESIL operates through a stack-based interpreter, akin to calculators, where values and operators are pushed and popped from a stack to perform operations. This post-fix notation allows for the expression of various CPU operations, including binary arithmetic, and memory operations. For example, Radare 2 can use ESIL to translate assembly instructions into understandable operations, shedding light on a program’s behavior, even for obscure architectures where traditional debugging tools may not be available. By harnessing ESIL’s capabilities, Radare 2 empowers analysts to efficiently identify syscall numbers and understand program execution flow.

Ghidra: How it does it?

Ghidra does not porovide an official way to solve the problem of guessing syscall numbers. Still, there’s a very well known third party script that uses Ghidra framework to provide the functionality. The Ghidra script employs symbolic propagation, a technique in reverse engineering that tracks symbolic values of program variables instead of executing the program. Specifically, the script analyzes x86 and x64 Linux binaries to resolve system calls. It identifies system call instructions, symbolically determines the values of syscall registers, and creates function definitions for each syscall. Symbolic propagation enables reasoning about register contents without actually executing the program.

My Proposition: Static Analysis + Emulation.

The initial concept is to emulate the code and examine the contents of registers when a syscall instruction is executed. However, a significant challenge arises: there's no assurance that the flow of code execution will necessarily reach the point of interest—the syscall.

To address this challenge, I employ a technique I refer to as guided execution. Essentially, this method involves guiding the execution to ensure it passes over the specific instruction I need to examine.

In the general case, this approach wouldn't provide guarantees about the accuracy of register contents. However, when it comes to syscalls, I am confident that the path I traverse to reach a syscall cannot determine the syscall number. This confidence stems from the fact that syscalls have diverse interfaces, implying that if multiple paths exist from the function start to a syscall, they will ultimately lead to the same syscall number. Subsequently, I outline the detailed procedure I employ to achieve this.

The process commences with the user-provided function binary image, a representation of compiled machine code specific to a given function. This binary image undergoes disassembly using the libcapstone disassembler, converting the machine code into a list of assembly instructions.

Following disassembly, an elaborate procedure ensues to generate the CFG for the function in question. The CFG serves as a graphical depiction of all feasible execution paths within the function. Each node within the graph corresponds to an instruction block, with edges denoting potential transitions between these blocks.

An illustrative example of this intermediate step pertains to the glibc function __pthread_mutex_lock_full.

Within this resulting CFG image, purple blocks denote instruction blocks housing syscalls, while cyan bubbles signify blocks containing a ret instruction.

Subsequently, the task shifts to identifying a path from the function’s entry point to the instruction block housing the syscall.

Once the path is discerned, a further operation ensues, scanning the code to locate any call instructions along the path. These call instructions are then substituted with nop instructions, ensuring guided execution remains focused solely on the path leading to the syscall instruction block and doesn’t veer off due to calls to other functions absent from the binary image.

With call instructions replaced, guided execution is executed using libunicorn. During this process, each instruction block along the path is executed sequentially, with the context (CPU state) of each block utilized to initiate execution of the subsequent block.

Upon reaching the syscall instruction, register values are inspected to ascertain the syscall number.

Addressing concerns and outlining why it’s effective in this use case.

The PoC

To apply this concept in a practical context, I aimed to create a Proof of Concept (PoC), opting to develop it as a Radare2 plugin. Below, I present the implementation for both versions of Radare:

Radare2.

Rizin.

Patching the Code

As I delve into the intricacies of Linux binaries, the necessity of patching becomes evident in my quest to uncover the elusive syscall numbers. Patching serves as a crucial tool to remove distractions and ensure a clear path to my destination.

However, the task of patching is notably more complex in CISC architectures compared to RISC. In CISC, such as x86, branch instructions vary in length, adding layers of intricacy to the process. This complexity is compounded by the need to decipher the true meaning behind seemingly innocuous instructions like the nop instruction.

In the realm of x86 architecture, the nop instruction is a facade for the xchg (e)ax, (e)ax instruction. This revelation sheds light on the true nature of nop, enabling me to craft multibyte instructions that effectively clear my path of obstructions.

To achieve this, I employ the Gauss formula, a stroke of genius that allows me to create a function capable of generating multibyte nop instructions of any desired size. Armed with this solution, I navigate the intricate landscape of x86 architecture with newfound confidence, inching closer to my ultimate goal with each patch applied.

const char multibyte_nop_x86[]= "\x90" "\x66\x90" "\x0f\x1f\x00" "\x0f\x1f\x40\x00" "\x66\x66\x66\x66\x90" "\x66\x0f\x1f\x44\x00\x00" "\x0f\x1f\x80\x00\x00\x00\x00" "\x0f\x1f\x84\x00\x00\x00\x00\x00" "\x66\x0f\x1f\x84\x00\x00\x00\x00\x00"; const char multibyte_ret_nop_x86[]= "\xc3" "\xc3\x90" "\xc3\x66\x90" "\xc3\x0f\x1f\x00" "\xc3\x0f\x1f\x40\x00" "\xc3\x66\x66\x66\x66\x90" "\xc3\x66\x0f\x1f\x44\x00\x00" "\xc3\x0f\x1f\x80\x00\x00\x00\x00" "\xc3\x0f\x1f\x84\x00\x00\x00\x00\x00"; #define MBNOP(n,t) (t==NOP_ONLY?multibyte_nop_x86 + (((n-1) * n)/2):multibyte_ret_nop_x86 + (((n-1) * n)/2))

No Return Functions.

The primary method employed by this plugin is emulation, where it does not have access to a context, meaning that function arguments are undetermined. Instead of emulating the entire function, it selectively executes the sequence of blocks of code leading from the function’s entry point to the point where a syscall is invoked, which is the guided execution I presented earlier.

To achieve this, the plugin performs an analysis of the target function and creates a CFG. The CFG serves as a navigational tool to identify the path of blocks leading to the block that contains the syscall.

In this context, “blocks,” or instruction blocks, are defined as sequences of instructions without any branching, typically terminating at a branch instruction. CALL instructions are typically treated as NOP instructions and are patched out of the binary.

However, a unique challenge arises when dealing with the glibc library, specifically, when encountering the __libc_fatal function. Unlike other functions, this troublemaker does not conform to the assumed behavior of returning, rendering the standard treatment of call instructions inappropriate.

I’ve decided a special treatment for this __libc_fatal function implementing ad hoc solution that transforms it into a RET instruction, allowing the plugin to function effectively in such cases.

Jump Tables: Tackling complexities introduced by jump tables.

The current plugin implementation effectively follows the CFG of a given function, providing accurate results in most scenarios. However, there’s an edge case that the current algorithm cannot handle correctly. This occurs when a switch case is implemented using a jump table. The current approach handles jmp instructions by considering no forward path, relying solely on the branch path. However, in cases involving jump tables, the branch path can be multiple and located far away from the code under examination.

A temporary solution focusing specifically on the syscall number can be summarized as follows:

Collection of Syscall Numbers: During the examination of the function’s code, the plugin should collect information about syscall numbers and their corresponding positions within the code.

Handling Non-Reached CFG Blocks: If the CFG analysis encounters non-reached blocks containing a syscall, the plugin could execute them to determine if the syscall number can be deduced.

Epilogue:

This endeavor stands as a PoC for a concept, while this blog post acts as a platform for others to voice their thoughts on this notion. I find it quite promising, although it requires refinement on multiple fronts. Summarizing key insights and takeaways. Encouragement for further exploration and research in the realm of binary analysis.

Navigating the Labyrinth: Battling Duplicate Symbols in the Kernel

Introduction

Hey folks! Ever wondered what it’s like to venture into the wild world of kernel development? Well, let me tell you, it’s not a walk in the park. In this post, I want to tell you about when I stumbled upon a head-scratcher – the conundrum of duplicate symbols in the Linux kernel.

Why Duplicate Symbols Exist in the Kernel and How is it Possible?

Kernel development is a complex and intricate process that involves the compilation and linking of numerous C source files to create the final executable. The kernel comprises multiple C files, each translated into an object file during compilation. The linking process that follows, is where these object files are brought together for form intermediate objects (or modules) and at last the kernel binary image. The linker relies on symbols, identifiers that represent variables or functions, to establish connections between different parts of the code. Importantly, these symbols the linker needs to use, must be unique within a linking session to ensure the process success.

In any given object file, there are two types of symbols: exported and non-exported. Exported symbols, designated to be used externally, must be unique within the set of objects the needs to be linked together to ensure the linking process success. On the other hand, non-exported symbols, often associated with functions declared as static, are not subject to the same uniqueness constraints, and the object linked together may end up in containing non exported symbols names that are not unique within the set.

As said, exported symbols must be unique to permit the linker’s job, but no strict rules are imposed on non-exported symbols. This characteristic is not something that must exist in general, but it is inevitably true in the vast codebase of the Linux kernel, where this occurrence is not uncommon.

Adding complexity to this scenario is the use of header files. Some symbols, particularly small inline functions, find their definition in these header files. While this is a common practice, it introduces the possibility of duplicate symbols, and when this happens, it is not just the result of chance, it is something structural that produces duplicates scientifically, since header files can be included in different compilation units, therefore the same function may appear multiple times within the kernel executable.

A notable contributor to the presence of duplicated symbols lies in the compatibility binary format ELF, where a hack in the compat_binfmt_elf.c file introduces duplicate symbols as a consequence of c file inclusion.

In essence, the kernel’s symbols are not granted to be unique in any way, and in fact there are duplicated symbols in the kernel.

~ # cat /proc/kallsyms | grep " name_show" ffffcaa2bb4f01c8 t name_show ffffcaa2bb9c1a30 t name_show ffffcaa2bbac4754 t name_show ffffcaa2bbba4900 t name_show ffffcaa2bbec4038 t name_show ffffcaa2bbecc920 t name_show ffffcaa2bbed3840 t name_show ffffcaa2bbef7210 t name_show ffffcaa2bbf03328 t name_show ffffcaa2bbff6f3c t name_show ffffcaa2bbff8d78 t name_show ffffcaa2bbfff7a4 t name_show ffffcaa2bc001f60 t name_show ffffcaa2bc009890 t name_show ffffcaa2bc01212c t name_show ffffcaa2bc025e2c t name_show ffffcaa2a052102c t name_show [hello] ffffcaa2a051955c t name_show [rpmsg_char]

Ok, there are duplicate symbols, but does this really matters?

It is a fair question, and I can anticipate that you can live a long life full of joy without even knowing about this problem. You can be even a kernel developer at some level and ignore the thing. Some kernel maintenance people cannot ignore the issue, though, since it won’t spare them from the pain of facing it. Whoever has done kernel tracing at some level is not spared by stomping on it, and also the few people who did live patching had to deal with this at a harder level. To understand how kernel duplicated symbol names can impact anyone, we must start from the fundamentals of the tracing system, the kernel symbol table aka kallsyms. The kallsyms is a giant table where any kernel symbol is noted together with its address. This table allows the tracing subsystem to locate a function within the kernel address space with a simple lookup. All this works particularly well; from a name, you have an address, as long as you have unique symbols.

Tracing a duplicate symbol, on the other hand, is a pain since you cannot know if the address kallsyms provides is the one you intended to watch. Saying that you can try again in the tracing usecase, things become more serious if you need to patch it.

Consider the scenario of tracing a specific function for debugging or performance analysis. With duplicate symbols, the tracer picks the one of the matching symbols (typically the first), but nobody knows if it is tho one you intended, leading to inaccurate results and potentially misinforming the trace engineer. The challenge amplifies when it comes to live patching, a process that involves modifying the kernel code at runtime. In such a delicate operation, precision is paramount, and the presence of duplicate symbols introduces an element of uncertainty. Need to say that the livepatch subsystem has somehow solved the problem of the uncertainty by using kallsyms_on_each_match_symbol, but it is still a suboptimal solution.

In essence, the existence of duplicate symbols adds an extra layer of complexity for those involved in kernel tracing and live patching.

Does this Really Happen, or is it Just a Potential Problem?

Well, the first time I stumbled upon this issue was when I began developing a tool to generate diagrams illustrating the dependencies between functions within the Linux kernel. Enter ks-nav, a tool designed to create graphs for any Linux kernel function, showcasing potential relations between various kernel subsystems. The software analyzes the kernel, exporting information into a database used to produce insightful diagrams. During this endeavor, I encountered the duplicate symbols issue for the first time. For my use case, I simply needed to address the duplicated names by renaming them adding leading sequence numbers, a seemingly manageable task at the time.

However, my subsequent encounter with duplicate symbols proved to be more challenging, and I instinctively steered away from confronting it directly. I was assigned the task of investigating a bug in a system that experienced a complete freeze and became unusable when utilizing ftrace functions. Without delving into the intricate details of the bug, the crux of the matter was that a peripheral generated an excessive number of interrupts, overwhelming the system to the point where the ftrace overhead left no CPU time for other operations.

Having identified the root cause, my intention was to create a demonstrator that would vividly illustrate the source of the problem. I aimed to hard-code a filter into the interrupt controller code to exclude the troublesome interrupt from being serviced. My strategy involved live-patching a function into the interrupt controller code, replacing it with the same function but with the filter hard-coded. It was a nice strategy, except that live patch does not support aarch64, and the interrupt controller symbol I wanted to replace, existed in thr two interrupt controller drivers built in the target kernel: the GIC and the GICv3.

Here are the links to the relevant kernel source files:

• GIC Interrupt Controller Driver
• GICv3 Interrupt Controller Driver

I managed to find a workaround for the issue, albeit in a somewhat hacky manner. However, this experience prompted me to recognize the need for a more generalized solution to address the challenges posed by duplicate symbols in the kernel.

In Any Case, What Havoc Do These Duplicate Symbols Wreak?

Let’s assume that the kernel itself doesn’t exhibit any substantial behavioral defects, and this issue predominantly affects the tracing system. Having a name that potentially does not precisely identify a function raises three significant issues with the current implementation. Let’s delve into the details:

1. Uncertainty in Symbol Address: The function responsible for looking up a name in the Linux kernel is kallsyms_lookup_name. It returns the address of the first symbol it encounters with that name. This introduces a problem because, given a symbol name, you can’t be certain that the address you receive corresponds to the specific function you’re interested in.
2. Identification Challenge: With multiple matching symbols, it’s not straightforward to determine which one is the intended symbol with a given name. The lack of a clear distinction complicates the selection process.
3. Difficulty in Addressing Specific Symbols: Assuming you’re aware that the first matching symbol with a given name is not the one you seek, it becomes challenging to address any other symbol. The process of singling out the correct symbol becomes nontrivial.

To address these challenges, the kernel provides kallsyms_lookup_names, which returns all symbols with the same name. While this can help refine the search, it doesn’t entirely resolve the issue. The live patch subsystem uses kallsyms_lookup_names in conjunction with kallsyms_on_each_match_symbol to apply a function on each entry and select the appropriate one. However, this approach has limitations, and its application is primarily confined to the live patch subsystem. If the function to apply as a filter were a BPF, it would be relatively more manageable.

The underlying problem is that, even with these mechanisms, kernel developers can only select the desired symbol in a convoluted way. From a user-space perspective, a trace engineer is left with addresses that are not particularly user-friendly. Moreover, determining the nature of a symbol is solely reliant on the address context, requiring the trace engineer to make educated guesses based on the available information. This complexity adds an extra layer of difficulty for those working with kernel tracing, underscoring the impact of duplicate symbols on practical kernel development scenarios.

So, is there any solution to fix this?

Before delving into solutions, it’s essential to acknowledge additional requirements that any potential resolution should meet:

4. Preserve the current state to not impact the ecosystem: Any solution addressing duplicated symbols in the kernel should strive to maintain the current state, ensuring minimal disruption to the ecosystem or any existing solutions implemented by trace software.

Now, let’s explore the possible solutions. A straightforward approach involves renaming all duplicate symbols to make each symbol unique. However, this is no trivial task, given the significant changes required. Moreover, there’s currently no preventive mechanism in place to avoid the recurrence of this issue. Any attempt to pursue this path should include a strategy to prevent the kernel from encountering the same situation in the future. It’s worth noting that renaming several symbols simultaneously could have unforeseeable impacts on the ecosystem.

So, what other options are available to address this issue? As of my investigation in July 2023, there were only two attempts to tackle the “Identification Challenge.”

My proposal to address this problem involves adding aliases to duplicate symbols.

Adding Aliases to Duplicate Symbols, can be really solve the problem?

My approach to resolving this problem is to introduce aliases for duplicate symbols. I believe it can provide a valuable solution for the problem, and here’s how aliases can address the identified issues:

1. Uncertainty in Symbol Address: Aliases need to be unique. Tracing a unique alias allows users to ensure that the provided address corresponds to the desired function.
2. Identification Challenge: The alias embodies information useful for correctly understanding the symbol’s nature or source.
3. Difficulty in Addressing Specific Symbols: The uniqueness of the alias allows the symbol to be selected simply by using the alias.
4. Preserve the Current State to Not Impact the Ecosystem: By adding aliases without removing the previous symbols, any ecosystem software that dealt with the problem in its way is preserved. This approach allows for a gradual migration to using aliases.

Indeed, special attention must be given to how aliases are generated. After consulting with the community, I opted to use a string as the alias, comprising the file path and the line number where the symbol is defined. This ensures the alias is not only unique but also carries crucial information for identifying the symbol.

Is yours proposal something concrete, or is it just words in the air?

This proposal is tangible and has undergone several iterations, currently standing at its 7th version as of January 2024. Detailed implementation discussions are more suited for the mailing list or the GitHub repository, where the development is actively taking place.

In its current state, the patch can:

• Automatically detect symbols requiring aliases during the kernel build.
• Add aliases to the core kernel image.
• Add aliases to modules.
• Export a file with symbol name statistics.
• Add aliases to out-of-tree kernel modules using symbol statistics.

This is how the symbols should look with the patch applied:

~ # cat /proc/kallsyms | grep " name_show" ffffcaa2bb4f01c8 t name_show ffffcaa2bb4f01c8 t name_show@kernel_irq_irqdesc_c_264 ffffcaa2bb9c1a30 t name_show ffffcaa2bb9c1a30 t name_show@drivers_pnp_card_c_186 ffffcaa2bbac4754 t name_show ffffcaa2bbac4754 t name_show@drivers_regulator_core_c_678 ffffcaa2bbba4900 t name_show ffffcaa2bbba4900 t name_show@drivers_base_power_wakeup_stats_c_93 ffffcaa2bbec4038 t name_show ffffcaa2bbec4038 t name_show@drivers_rtc_sysfs_c_26 ffffcaa2bbecc920 t name_show ffffcaa2bbecc920 t name_show@drivers_i2c_i2c_core_base_c_660 ffffcaa2bbed3840 t name_show ffffcaa2bbed3840 t name_show@drivers_i2c_i2c_dev_c_100 ffffcaa2bbef7210 t name_show ffffcaa2bbef7210 t name_show@drivers_pps_sysfs_c_66 ffffcaa2bbf03328 t name_show ffffcaa2bbf03328 t name_show@drivers_hwmon_hwmon_c_72 ffffcaa2bbff6f3c t name_show ffffcaa2bbff6f3c t name_show@drivers_remoteproc_remoteproc_sysfs_c_215 ffffcaa2bbff8d78 t name_show ffffcaa2bbff8d78 t name_show@drivers_rpmsg_rpmsg_core_c_455 ffffcaa2bbfff7a4 t name_show ffffcaa2bbfff7a4 t name_show@drivers_devfreq_devfreq_c_1395 ffffcaa2bc001f60 t name_show ffffcaa2bc001f60 t name_show@drivers_extcon_extcon_c_389 ffffcaa2bc009890 t name_show ffffcaa2bc009890 t name_show@drivers_iio_industrialio_core_c_1396 ffffcaa2bc01212c t name_show ffffcaa2bc01212c t name_show@drivers_iio_industrialio_trigger_c_51 ffffcaa2bc025e2c t name_show ffffcaa2bc025e2c t name_show@drivers_fpga_fpga_mgr_c_618 ffffcaa2a052102c t name_show [hello] ffffcaa2a052102c t name_show@hello_hello_c_8 [hello] ffffcaa2a051955c t name_show [rpmsg_char] ffffcaa2a051955c t name_show@drivers_rpmsg_rpmsg_char_c_365 [rpmsg_char]

However, refinement and alignment with community standards are still ongoing, requiring further iterations. The ultimate acceptance of this work into the kernel remains uncertain, but this article serves as a testament to the effort invested.

Unveiling CPU Count in a System: A Dive into Linux Utilities and Functions

Introduction:

Understanding the computational capabilities of a system can be pivotal in optimizing software performance. The quest to determine the number of CPUs within a system often leads to exploration and evaluation of various methods. This pursuit becomes particularly critical when developing low-level tools for Linux, where having an accurate insight into the available CPU count is essential for efficient resource utilization and task allocation. In this article, I delve into different approaches to uncovering this crucial system attribute, aiming to provide insights for developers grappling with similar challenges.

Initial Proposition:

In my pursuit of determining the number of CPUs within a Linux system, my initial approach mirrored what I typically did from the command line: cat /proc/cpuinfo. It’s worth noting that while the command line utility nproc caters to the same query, my familiarity with the proc method predates my acquaintance with nproc. Hence, it was my instinctive choice for the initial implementation of a function to retrieve this vital information programmatically.

int nproc() {
        int cpu_count = 0;
        char line[100];
        FILE *file;

        file = fopen("/proc/cpuinfo", "r");
        if (file == NULL) {
                perror("Error opening file");
                exit(EXIT_FAILURE);
        }

        while (fgets(line, sizeof(line), file)) {
                if (strncmp(line, "processor", 9) == 0) {
                        cpu_count++;
                }
        }

        fclose(file);
        return cpu_count;
}

Does this implementation work? Yes, it does provide the expected CPU count. However, despite its functionality, I find myself dissatisfied with this method for several reasons, which I’ll delve into shortly.

Concerns on the Solution:

While the initial /proc/cpuinfo approach effectively retrieves CPU information, several concerns arise, prompting dissatisfaction with this simplistic method:

1. Oversimplicity: The code’s apparent simplicity, merely opening a file and parsing statements, raises skepticism. Its straightforward nature seems inadequate for a task as critical as determining CPU count within a system. This simplicity might overlook nuances or potential intricacies within different system environments.
2. Uncertainty in /proc/cpuinfo Stability: /proc/cpuinfo provides information designed for human readability. This readability doesn’t guarantee immutable structure or content, leaving room for future modifications by kernel developers. While changes aren’t imminent, relying solely on this file may pose a risk of unexpected alterations in its format or content, impacting the function’s reliability.
3. Dependency on Procfs: The reliance on procfs introduces a constraint, assuming its presence and accessibility, which isn’t universally guaranteed. In specific embedded Linux systems or constrained environments, procfs might not be mounted or accessible. Assumptions like these could jeopardize the function’s portability and reliability across diverse Linux distributions and setups.
4. Preference for Syscall Stability: A preference emerges for utilizing syscalls, known for their stability and less prone to alterations compared to file-based interfaces like procfs. Syscalls, by design, tend to remain more consistent across Linux distributions and versions, offering a more robust foundation for critical functionalities like retrieving CPU count.

The sysconf() Function:

Is there a standardized, reliable way to retrieve CPU count programmatically? After a comprehensive search, one function stood out as a prevailing best practice: sysconf(_SC_NPROCESSORS_CONF). This function, available in the POSIX standard, offers a promising solution for obtaining the number of processors in a system.

What is sysconf()?

At its core, sysconf() is a system call that allows access to configurable system variables. Specifically, _SC_NPROCESSORS_CONF is a parameter used to query the number of processors configured in the system. Additionally, _SC_NPROCESSORS_ONLN exists, reporting the number of CPUs effectively active.

Does it provide a solution for this?

Indeed, sysconf(_SC_NPROCESSORS_CONF) serves as a robust and standardized means to obtain CPU count programmatically. Its utilization ensures compatibility across various Linux distributions and versions, contributing to code portability.

How sysconf() is implemented in the guts of libcs:

glibc implementation:

Navigating through the labyrinthine structure of glibc, a colossal library accommodating various systems, one discovers numerous implementations catering to different operating environments. Amidst this complexity, glibc’s handling of sysconf() for Linux systems proves enlightening.

In scrutinizing glibc’s codebase, particularly for Linux-specific implementations, it’s evident that glibc relies on sysfs as its primary source for retrieving CPU count information. The main method employed by glibc to provide this data involves accessing /sys/devices/system/cpu/possible. This file seemingly contains the requisite CPU count information, hence serving as the cornerstone for glibc’s approach.

However, acknowledging the potential unpredictability of specialized file system resources, glibc developers have implemented a fallback system in case the primary method fails. The fallback function, get_nprocs_fallback(), comprises alternative methods to retrieve CPU count information:

1. get_nproc_stat(): This method mirrors my initial implementation; it parses /proc/cpuinfo and tallies occurrences of cpu, akin to counting processors. While functional, it remains subject to the same concerns surrounding /proc/cpuinfo stability and dependencies on procfs availability.

2. __get_nprocs_sched(): Addressing concerns about specialized filesystem accessibility, this method adopts sched_getaffinity(), essentially a syscall, to deduce the number of CPUs. This syscall-based approach aligns with the desire for a more reliable, system-level means to retrieve CPU count information, bypassing potential dependencies on file system structures.

Despite glibc’s vastness and varied system support, its Linux implementation of sysconf() prioritizes sysfs as the primary resource for CPU count retrieval, supplemented by fallback methods for contingencies or specialized system scenarios.

musl implementation:

Developed with a distinct focus on constrained systems, musl libc embodies the quintessential choice for environments with stringent resource limitations. Unveiling musl’s methodology for retrieving CPU count information unveils a strategy tailored to suit such constrained environments.

In contrast to glibc’s expansive versatility, musl prioritizes efficiency and minimalism, aligning with the requirements of resource-constrained systems. musl’s approach to CPU count retrieval centers around a singular method, reliant on the sched_getaffinity() syscall to infer CPU count.

This method, employing sched_getaffinity(), emphasizes a syscall-based approach for determining CPU count information. While not leveraging sysfs or other file system structures, musl’s implementation demonstrates a steadfast reliance on system-level calls for this critical system attribute. This approach aligns with musl’s ethos of simplicity and efficiency, avoiding potential dependencies on file system structures and instead relying solely on syscall interactions to obtain CPU count details.

musl’s singular reliance on sched_getaffinity() for CPU count retrieval underscores its commitment to efficiency and simplicity in constrained system environments, reflecting its status as a libc tailored specifically for resource-limited setups.

ulibc implementation:

ulibc, akin to musl, stands as a libc tailored explicitly for resource-constrained systems. Its implementation strategy for CPU count retrieval reflects a minimalistic approach, aligning with the requirements of such constrained environments.

Despite the expectation of multiple methods to cater to contingencies, ulibc’s decision to rely solely on a sysfs based approach might initially surprise. The implementation involves scanning the sysfs directory /sys/devices/system/cpu, specifically enumerating directory entries featuring the substring ‘cpu[0-9]’. This singular method utilizing sysfs as the source for CPU count determination aligns with ulibc’s ethos of minimalism and efficiency.

This reliance on sysfs, while potentially limiting if sysfs isn’t mounted, resonates with ulibc’s targeted service provision to highly constrained systems. The decision potentially avoids additional overhead that could arise from using syscall-based methods, as it requires a support structure having a non-negligible footprint, incompatible with resource-limited environments.

nproc Utility:

Examining the nproc utility, it’s pertinent to delve into two notable implementations: Coreutils and Busybox. Despite the varying nature of these utilities, the expectation leans toward both implementations relying on support from the underlying libc to provide CPU count information.

Coreutils Implementation:

Coreutils, a fundamental suite of Unix utilities including nproc, typically relies on libc support for system-specific information retrieval. The nproc utility in Coreutils is expected to utilize system calls or libc functions to fetch CPU count data. Upon inspecting the nproc utility’s codebase, it becomes evident that the utility itself does not house specific code for CPU count retrieval. Instead, it relies on certain environment variables typically provided by the OpenMP library. In the absence of these variables, nproc falls back on a function named num_processors_ignoring_omp(), which finds its implementation within the gnulib.

The gnulib implementation might initially appear complex, owing to its handling of the task across diverse platforms, including one for Windows. However, restricting the investigation to Linux unveils an implementation primarily reliant on the sched_getaffinity() syscall. Remarkably, this implementation appears independent of direct dependencies on the libc, offering an alternative method for CPU count retrieval.

By leaning on sched_getaffinity(), the gnulib implementation for Linux within nproc exemplifies a straightforward syscall-based approach for CPU count determination, suggesting a certain level of autonomy from standard libc functionalities.

Busybox Implementation:

Busybox, renowned for its compactness and versatility, reimagines system utilities, often independent of direct dependencies on standard libraries like libc. The implementation of CPU count retrieval within Busybox resonates with this independent ethos, showcasing a self-contained logic mirroring methodologies akin to libc implementations.

Busybox’s CPU count retrieval method primarily involves scanning the sysfs directory /sys/devices/system/cpu, specifically enumerating directory items containing the substring cpu. However, this method assumes prominence only when a configuration symbol is enabled, signifying Busybox’s adaptability based on configuration choices.

The surprising element lies in Busybox’s delineation of preferences: designating sysfs-based enumeration as the primary and fallback method, whereas the sched_getaffinity() syscall serves as an alternative choice. This nuanced approach differs from conventional expectations, underscoring Busybox’s nuanced perspective on reliability and system adaptability.

Despite sharing similarities with libc-based implementations in sysfs enumeration, Busybox’s developers view the syscall-based approach as a secondary yet dependable method. This acknowledgment highlights sched_getaffinity() as an alternative in scenarios permitting multiple methods, but also as the preferred method when only one can be chosen.

Consideration about Portability:

Having explored various existing methods for accessing CPU count data, it’s apparent that despite the availability of libc-provided methods, many userspace utilities opt to implement their own functions for this purpose. Reflecting on this, in a new implementation, the recommendation leans toward leveraging the libc-provided method, specifically sysconf(_SC_NPROCESSORS_CONF). This approach ensures compatibility and adherence to standard interfaces across diverse Linux environments.

However, in specific scenarios where unique constraints exist, such as employing ulibc in an environment without accessible or mountable special filesystems, it may necessitate a custom implementation for CPU count retrieval. In such cases, an alternative suggestion would be to base the implementation on the sched_getaffinity() syscall, given its usability across varied contexts.

Method	description	usable	memory footprint
Sysfs file /sys/devices/system/cpu/possible	the file gets accessed and it contains the number we need	when sysfs is mounted	lowest
sysfs directory /sys/devices/system/cpu	the directory is scanned and cpu[0-9]+ subdirectories are searched. Count the items provides the number.	when sysfs is mounted	low
procfs file parsing /proc/cpuinfo	the file is accessed and the contents parsed. Counting CPU occurence provides the number.	when procfs is mounted	medium
sched_getaffinity()	the scheduler affinity table is copied to the userspace, items are counted.	always	possibly high

sched_getaffinity() sample implementation

To end this article I want to provide a sched_getaffinity() based sample implementation.

#define _GNU_SOURCE
#include <sched.h>

#define MAX_CPUS 2048

int nproc(void){
        unsigned long mask[MAX_CPUS];
        unsigned long m;
        int count = 0;
        int i;

        if (sched_getaffinity(0, sizeof(mask), (void*)mask) == 0) {
                for (i = 0; i < MAX_CPUS; i++) {
                        m = mask[i];
                        while (m) {
                                if (m & 1) count++;
                                m >>= 1;
                        }
                }
        }
        return count;
}