Registers are the fastest kind of memory available in the machine. x86-64 has 14 general-purpose registers and several special-purpose registers. This table gives all the basic registers, with special-purpose registers highlighted in yellow. You’ll notice different naming conventions, a side effect of the long history of the x86 architecture (the 8086 was first released in 1978).

Full register
(bits 0-63)

(bits 0–31)

(bits 0–15)

8-bit low
(bits 0–7)

8-bit high
(bits 8–15)

Use in calling convention


General-purpose registers:






Return value (accumulator)













4th function argument







3rd function argument






2nd function argument






1st function argument






5th function argument






6th function argument
































Special-purpose registers:





Stack pointer






Base pointer
(general-purpose in some compiler modes)





Instruction pointer
(Program counter; called $pc in GDB)





Flags and condition codes


Note that unlike primary memory (which is what we think of when we discuss memory in a C/C++ program), registers have no addresses! There is no address value that, if cast to a pointer and dereferenced, would return the contents of the %rax register. Registers live in a separate world from the memory whose contents are partially prescribed by the C abstract machine.

The %rbp register has a special purpose: it points to the bottom of the current function’s stack frame, and local variables are often accessed relative to its value. However, when optimization is on, the compiler may determine that all local variables can be stored in registers. This frees up %rbp for use as another general-purpose register.

The relationship between different register bit widths is a little weird.

  1. Modifying a 32-bit register name sets the upper 32 bits of the register to zero. Thus, after movl $-1, %eax, the %rax register has value 0x0000'0000'FFFF'FFFF. The same is true after movq $-1, %rax; addl $0, %eax! (The movq sets %rax to 0xFFFF'FFFF'FFFF'FFFF; the addl sets its upper 32 bits to zero.)

  2. Modifying a 16- or 8-bit register name leaves all other bits of the register unchanged.

There are special instructions for loading signed and unsigned 8-, 16-, and 32-bit quantities into registers, recognizable by instruction suffixes. For instance, movzbl moves an 8-bit quantity (a byte) into a 32-bit register (a longword) with zero extension; movslq moves a 32-bit quantity (longword) into a 64-bit register (quadword) with sign extension. There’s no need for movzlq (why?).

Instruction format

The basic kinds of assembly instructions are:

  1. Arithmetic. These instructions perform computation on values, typically values stored in registers. Most have zero or one source operands and one source/destination operand. The source operand is listed first in the instruction, but the source/destination operand comes first in the computation (this matters for non-commutative operators like subtraction). For example, the instruction addq %rax, %rbx performs the computation %rbx := %rbx + %rax.

  2. Data movement. These instructions move data between registers and memory. Almost all have one source operand and one destination operand; the source operand comes first.

  3. Control flow. Normally the CPU executes instructions in sequence. Control flow instructions change the instruction pointer in other ways. There are unconditional branches (the instruction pointer is set to a new value), conditional branches (the instruction pointer is set to a new value if a condition is true), and function call and return instructions.

(We use the “AT&T syntax” for x86-64 assembly. For the “Intel syntax,” which you can find in online documentation from Intel, see the Aside in CS:APP3e §3.3, p177, or Wikipedia, or other online resources. AT&T syntax is distinguished by several features, but especially by the use of percent signs for registers. Sadly, the Intel syntax puts destination registers before source registers.)

Some instructions appear to combine arithmetic and data movement. For example, given the C code int* ip; ... ++(*ip); the compiler might generate incl (%rax) rather than movl (%rax), %ebx; incl %ebx; movl %ebx, (%rax). However, the processor actually divides these complex instructions into tiny, simpler, invisible instructions called microcode, because the simpler instructions can be made to execute faster. The complex incl instruction actually runs in three phases: data movement, then arithmetic, then data movement. This matters when we introduce parallelism.


Assembly generated by a compiler contains instructions as well as labels and directives. Labels look like labelname: or labelnumber:; directives look like .directivename arguments. Labels are markers in the generated assembly, used to compute addresses. We usually see them used in control flow instructions, as in jmp L3 (“jump to L3”). Directives are instructions to the assembler; for instance, the .globl L instruction says “label L is globally visible in the executable”, .align sets the alignment of the following data, .long puts a number in the output, and .text and .data define the current segment.

We also frequently look at assembly that is disassembled from executable instructions by GDB, objdump -d, or objdump -S. This output looks different from compiler-generated assembly: in disassembled instructions, there are no intermediate labels or directives. This is because the labels and directives disappear during the process of generating executable instructions.

For instance, here is some compiler-generated assembly:

    .globl  _Z1fiii
    .type   _Z1fiii, @function
    cmpl    %edx, %esi
    je  .L3
    movl    %esi, %eax
    movl    %edi, %eax
    .size   _Z1fiii, .-_Z1fiii

And a disassembly of the same function, from an object file:

0000000000000000 <_Z1fiii>:
   0:   39 d6                   cmp    %edx,%esi
   2:   74 03                   je     7 <_Z1fiii+0x7>
   4:   89 f0                   mov    %esi,%eax
   6:   c3                      retq   
   7:   89 f8                   mov    %edi,%eax
   9:   c3                      retq   

Everything but the instructions is removed, and the helpful .L3 label has been replaced with an actual address. The function appears to be located at address 0. This is just a placeholder; the final address is assigned by the linking process, when a final executable is created.

Finally, here is some disassembly from an executable:

0000000000400517 <_Z1fiii>:
  400517:   39 d6                   cmp    %edx,%esi
  400519:   74 03                   je     40051e <_Z1fiii+0x7>
  40051b:   89 f0                   mov    %esi,%eax
  40051d:   c3                      retq   
  40051e:   89 f8                   mov    %edi,%eax
  400520:   c3                      retq   

The instructions are the same, but the addresses are different. (Other compiler flags would generate different addresses.)

Address modes

Most instruction operands use the following syntax for values. (See also CS:APP3e Figure 3.3 in §3.4.1, p181.)

Type Example syntax Value used
Register %rbp Contents of %rbp
Immediate $0x4 0x4
Memory 0x4 Value stored at address 0x4
symbol_name Value stored in global symbol_name
(the compiler resolves the symbol name to an address when creating the executable)
symbol_name(%rip) %rip-relative addressing for global
symbol_name+4(%rip) Simple arithmetic on symbols are allowed
(the compiler resolves the arithmetic when creating the executable)
(%rax) Value stored at address in %rax
0x4(%rax) Value stored at address %rax + 4
(%rax,%rbx) Value stored at address %rax + %rbx
(%rax,%rbx,4) Value stored at address %rax + %rbx*4
0x18(%rax,%rbx,4) Value stored at address %rax + 0x18 + %rbx*4

The full form of a memory operand is offset(base,index,scale), which refers to the address offset + base + index*scale. In 0x18(%rax,%rbx,4), %rax is the base, 0x18 the offset, %rbx the index, and 4 the scale. The offset (if used) must be a constant and the base and index (if used) must be registers; the scale must be either 1, 2, 4, or 8. The default offset, base, and index are 0, and the default scale is 1.

symbol_names are found only in compiler-generated assembly; disassembly uses raw addresses (0x601030) or %rip-relative offsets (0x200bf2(%rip)).

Jumps and function call instructions use different syntax 🤷🏽‍♀️: *, rather than (), represents indirection.

Type Example syntax Address used
Register *%rax Contents of %rax
Immediate .L3 Address of .L3 (compiler-generated assembly)
400410 or 0x400410 Given address
Memory *0x200b96(%rip) Value stored at address %rip + 0x200b96
*(%r12,%rbp,8) Other address modes accepted

Address arithmetic

The base(offset,index,scale) form compactly expresses many array-style address computations. It’s typically used with a mov-type instruction to dereference memory. However, the compiler can use that form to compute addresses, thanks to the lea (Load Effective Address) instruction.

For instance, in movl 0x18(%rax,%rbx,4), %ecx, the address %rax + 0x18 + %rbx*4 is computed, then immediately dereferenced: the 4-byte value located there is loaded into %ecx. In leaq 0x18(%rax,%rbx,4), %rcx, the same address is computed, but it is not dereferenced. Instead, the computed address is moved into register %rcx.

Thanks to lea, the compiler will also prefer the base(offset,index,scale) form over add and mov for certain arithmetic computations on integers. For example, this instruction:

leaq (%rax,%rbx,2), %rcx

performs the function %rcx := %rax + 2 * %rbx, but in one instruction, rather than three (movq %rax, %rcx; addq %rbx, %rcx; addq %rbx, %rcx).

%rip-relative addressing

x86-64 code often refers to globals using %rip-relative addressing: a global variable named a is referenced as a(%rip) rather than a.

This style of reference supports position-independent code (PIC), a security feature. It specifically supports position-independent executables (PIEs), which are programs that work independently of where their code is loaded into memory.

To run a conventional program, the operating system loads the program’s instructions into memory at a fixed address that’s the same every time, then starts executing the program at its first instruction. This works great, and runs the program in a predictable execution environment (the addresses of functions and global variables are the same every time). Unfortunately, the very predictability of this environment makes the program easier to attack.

In a position-independent executable, the operating system loads the program at varying locations: every time it runs, the program’s functions and global variables have different addresses. This makes the program harder to attack (though not impossible).

Program startup performance matters, so the operating system doesn’t recompile the program with different addresses each time. Instead, the compiler does most of the work in advance by using relative addressing.

When the operating system loads a PIE, it picks a starting point and loads all instructions and globals relative to that starting point. The PIE’s instructions never refer to global variables using direct addressing: you’ll never see movl global_int, %eax. Globals are referenced relatively instead, using deltas relative to the next %rip: movl global_int(%rip), %eax. These relative addresses work great independent of starting point! For instance, consider an instruction located at (starting-point + 0x80) that loads a variable g located at (starting-point + 0x1000) into %rax. In a non-PIE, the instruction might be written movq 0x400080, %rax (in compiler output, movq g, %rax); but this relies on g having a fixed address. In a PIE, the instruction might be written movq 0xf79(%rip), %rax (in compiler output, movq g(%rip), %rax), which works out beautifully no matter the starting point.

At starting point… The mov instruction is at… The next instruction is at… And g is at… So the delta (g - next %rip) is…
0x400000 0x400080 0x400087 0x401000 0xF79
0x404000 0x404080 0x404087 0x405000 0xF79
0x4003F0 0x400470 0x400477 0x4013F0 0xF79

Arithmetic instructions

The operations of many x86-64 arithmetic instructions are easy enough to guess from their names. Then there are some arithmetic instructions, particularly those associated with Streaming SIMD Extensions and its follow-ons, that are hard to guess (phminposuw?).

The basic arithmetic instructions on 64-bit quantities (“quadwords”) are:

Instruction Operation Type Expansion
addq SRC, DST Addition Normal DST := DST + SRC
subq SRC, DST Subtraction Normal DST := DST - SRC
incq DST Increment Normal DST := DST + 1
decq DST Decrement Normal DST := DST - 1
imulq SRC, DST Signed multiplication Normal DST := DST * SRC
negq DST Negation Normal DST := -DST (DST := ~DST + 1)
andq SRC, DST Bitwise and Normal DST := DST & SRC
orq SRC, DST Bitwise or Normal DST := DST | SRC
xorq SRC, DST Bitwise exclusive or Normal DST := DST ^ SRC
notq DST Complement Normal DST := ~DST
sal SRC, DST
(also shl SRC, DST)
Left shift Normal DST := DST << SRC
sar SRC, DST Signed right shift Normal DST := DST >> SRC, shifting in sign bit
shr SRC, DST Unsigned right shift Normal DST := DST >> SRC, shifting in zeros
cmpq SRC, DST Subtraction for flags Flags-only DST - SRC; see below
testq SRC, DST Bitwise and for flags Flags-only DST & SRC; see below

There are also compact multiplication and division instructions that modify multiple registers at once and take fixed registers for some arguments. These instructions treat the combination of %rax and %rdx as a single 128-bit value where the most significant bits (bits 64–127) of the value are stored in %rdx and the least significant bits (0–63) are stored in %rax. The division instructions compute both a quotient and a remainder. (In the below, TMP is a 128-bit number.)

Instruction Operation Type Expansion
imulq SRC Signed multiplication Mul/div TMP := %rax * SRC; %rdx := TMP>>64; %rax := TMP
mulq SRC Unsigned multiplication Mul/div TMP := %rax * SRC; %rdx := TMP>>64; %rax := TMP
idivq SRC Signed division Mul/div TMP := (%rdx<<64) | %rax; %rax := TMP / SRC; %rdx := TMP % SRC
divq SRC Unsigned division Mul/div TMP := (%rdx<<64) | %rax; %rax := TMP / SRC; %rdx := TMP % SRC

Calling convention

A calling convention governs how functions on a particular architecture and operating system interact. This includes rules about includes how function arguments are placed, where return values go, what registers functions may use, how they may allocate local variables, and so forth. Calling conventions ensure that functions compiled by different compilers can interoperate, and they ensure that operating systems can run code from different programming languages and compilers. Some aspects of a calling convention are derived from the instruction set itself, but some are conventional, meaning decided upon by people (for instance, at a convention).

Calling conventions constrain both callers and callees. A caller is a function that calls another function; a callee is a function that was called. The currently-executing function is a callee, but not a caller.

For concreteness, we learn the x86-64 calling conventions for Linux. These conventions are shared by many OSes, including MacOS (but not Windows), and are officially called the “System V AMD64 ABI.”

The official specification: AMD64 ABI

Argument passing and stack frames

One set of calling convention rules governs how function arguments and return values are passed. On x86-64 Linux, the first six function arguments are passed in registers %rdi, %rsi, %rdx, %rcx, %r8, and %r9, respectively. The seventh and subsequent arguments are passed on the stack, about which more below. The return value is passed in register %rax.

The full rules more complex than this. You can read them in the AMD64 ABI, section 3.2.3, but they’re quite detailed. Some highlights:

  1. A structure argument that fits in a single machine word (64 bits/8 bytes) is passed in a single register.

    Example: struct small { char a1, a2; }

  2. A structure that fits in two to four machine words (16–32 bytes) is passed in sequential registers, as if it were multiple arguments.

    Example: struct medium { long a1, a2; }

  3. A structure that’s larger than four machine words is always passed on the stack.

    Example: struct large { long a, b, c, d, e, f, g; }

  4. Floating point arguments are generally passed in special registers, the “SSE registers,” that we don’t discuss further.

  5. Return values of up to one machine word are passed in %rax. Many return values that fit in two machine words—for instance, a pair of longs—are passed in %rax (for the first 8 bytes) and %rdx (for the second 8 bytes). For return values that fit only in three or more machine words, the caller reserves space for the return value, and passes the address of that space as the first argument of the function. The callee will fill in that space when it returns.

Writing small programs to demonstrate these rules is a pleasant exercise; for example:

struct small { char a1, a2; };
int f(small s) {
    return s.a1 + 2 * s.a2;

compiles to:

movl %edi, %eax           # copy argument to %eax
movsbl %dil, %edi         # %edi := sign-extension of lowest byte of argument (s.a1)
movsbl %ah, %eax          # %eax := sign-extension of 2nd byte of argument (s.a2)
movsbl %al, %eax
leal (%rdi,%rax,2), %eax  # %eax := %edi + 2 * %eax


Recall that the stack is a segment of memory used to store objects with automatic lifetime. Typical stack addresses on x86-64 look like 0x7ffd'9f10'4f58—that is, close to 247.

The stack is named after a data structure, which was sort of named after pancakes. Stack data structures support at least three operations: push adds a new element to the “top” of the stack; pop removes the top element, showing whatever was underneath; and top accesses the top element. Note what’s missing: the data structure does not allow access to elements other than the top. (Which is sort of how stacks of pancakes work.) This restriction can speed up stack implementations.

Like a stack data structure, the stack memory segment is only accessed from the top. The currently running function accesses its local variables; the function’s caller, grand-caller, great-grand-caller, and so forth are dormant until the currently running function returns.

x86-64 stacks look like this:


The x86-64 %rsp register is a special-purpose register that defines the current “stack pointer.” This holds the address of the current top of the stack. On x86-64, as on many architectures, stacks grow down: a “push” operation adds space for more automatic-lifetime objects by moving the stack pointer left, to a numerically-smaller address, and a “pop” operation recycles space by moving the stack pointer right, to a numerically-larger address. This means that, considered numerically, the “top” of the stack has a smaller address than the “bottom.”

This is built in to the architecture by the operation of instructions like pushq, popq, call, and ret. A push instruction pushes a value onto the stack. This both modifies the stack pointer (making it smaller) and modifies the stack segment (by moving data there). For instance, the instruction pushq X means:

subq $8, %rsp
movq X, (%rsp)

And popq X undoes the effect of pushq X. It means:

movq (%rsp), X
addq $8, %rsp

X can be a register or a memory reference.

The portion of the stack reserved for a function is called that function’s stack frame. Stack frames are aligned: x86-64 requires that each stack frame be a multiple of 16 bytes, and when a callq instruction begins execution, the %rsp register must be 16-byte aligned. This means that every function’s entry %rsp address will be 8 bytes off a multiple of 16.

Return address and entry and exit sequence

The steps required to call a function are sometimes called the entry sequence and the steps required to return are called the exit sequence. Both caller and callee have responsibilities in each sequence.

To prepare for a function call, the caller performs the following tasks in its entry sequence.

  1. The caller stores the first six arguments in the corresponding registers.

  2. If the callee takes more than six arguments, or if some of its arguments are large, the caller must store the surplus arguments on its stack frame. It stores these in increasing order, so that the 7th argument has a smaller address than the 8th argument, and so forth. The 7th argument must be stored at (%rsp) (that is, the top of the stack) when the caller executes its callq instruction.

  3. The caller saves any caller-saved registers (see below).

  4. The caller executes callq FUNCTION. This has an effect like pushq $NEXT_INSTRUCTION; jmp FUNCTION (or, equivalently, subq $8, %rsp; movq $NEXT_INSTRUCTION, (%rsp); jmp FUNCTION), where NEXT_INSTRUCTION is the address of the instruction immediately following callq.

This leaves a stack like this:

Initial stack at start of function

To return from a function:

  1. The callee places its return value in %rax.

  2. The callee restores the stack pointer to its value at entry (“entry %rsp”), if necessary.

  3. The callee executes the retq instruction. This has an effect like popq %rip, which removes the return address from the stack and jumps to that address.

  4. The caller then cleans up any space it prepared for arguments and restores caller-saved registers if necessary.

Particularly simple callees don’t need to do much more than return, but most callees will perform more tasks, such as allocating space for local variables and calling functions themselves.

Callee-saved registers and caller-saved registers

The calling convention gives callers and callees certain guarantees and responsibilities about the values of registers across function calls. Function implementations may expect these guarantees to hold, and must work to fulfill their responsibilities.

The most important responsibility is that certain registers’ values must be preserved across function calls. A callee may use these registers, but if it changes them, it must restore them to their original values before returning. These registers are called callee-saved registers. All other registers are caller-saved.

Callers can simply use callee-saved registers across function calls; in this sense they behave like C++ local variables. Caller-saved registers behave differently: if a caller wants to preserve the value of a caller-saved register across a function call, the caller must explicitly save it before the callq and restore it when the function resumes.

On x86-64 Linux, %rbp, %rbx, %r12, %r13, %r14, and %r15 are callee-saved, as (sort of) are %rsp and %rip. The other registers are caller-saved.

Base pointer (frame pointer)

The %rbp register is called the base pointer (and sometimes the frame pointer). For simple functions, an optimizing compiler generally treats this like any other callee-saved general-purpose register. However, for more complex functions, %rbp is used in a specific pattern that facilitates debugging. It works like this:

Stack frame with base pointer

  1. The first instruction executed on function entry is pushq %rbp. This saves the caller’s value for %rbp into the callee’s stack. (Since %rbp is callee-saved, the callee must save it.)

  2. The second instruction is movq %rsp, %rbp. This saves the current stack pointer in %rbp (so %rbp = entry %rsp - 8).

    This adjusted value of %rbp is the callee’s “frame pointer.” The callee will not change this value until it returns. The frame pointer provides a stable reference point for local variables and caller arguments. (Complex functions may need a stable reference point because they reserve varying amounts of space for calling different functions.)

    Note, also, that the value stored at (%rbp) is the caller’s %rbp, and the value stored at 8(%rbp) is the return address. This information can be used to trace backwards through callers’ stack frames by functions such as debuggers.

  3. The function ends with movq %rbp, %rsp; popq %rbp; retq, or, equivalently, leave; retq. This sequence restores the caller’s %rbp and entry %rsp before returning.

Stack size and red zone

Functions execute fast because allocating space within a function is simply a matter of decrementing %rsp. This is much cheaper than a call to malloc or new! But making this work takes a lot of machinery. We’ll see this in more detail later; but in brief: The operating system knows that %rsp points to the stack, so if a function accesses nonexistent memory near %rsp, the OS assumes it’s for the stack and transparently allocates new memory there.

So how can a program “run out of stack”? The operating system puts a limit on each function’s stack, and if %rsp gets too low, the program segmentation faults.

The diagram above also shows a nice feature of the x86-64 architecture, namely the red zone. This is a small area above the stack pointer (that is, at lower addresses than %rsp) that can be used by the currently-running function for local variables. The red zone is nice because it can be used without mucking around with the stack pointer; for small functions push and pop instructions end up taking time.


The processor typically executes instructions in sequence, incrementing %rip each time. Deviations from sequential instruction execution, such as function calls, are called control flow transfers.

Function calls aren’t the only kind of control flow transfer. A branch instruction jumps to a new instruction without saving a return address on the stack.

Branches come in two flavors, unconditional and conditional. The jmp or j instruction executes an unconditional branch (like a goto). All other branch instructions are conditional: they only branch if some condition holds. That condition is represented by condition flags that are set as a side effect of every arithmetic operation.

Arithmetic instructions change part of the %rflags register as a side effect of their operation. The most often used flags are:

Flags are most often accessed via conditional jump or conditional move instructions. The conditional branch instructions are:

Instruction Mnemonic C example Flags
j (jmp) Jump break; (Unconditional)
je (jz) Jump if equal (zero) if (x == y) ZF
jne (jnz) Jump if not equal (nonzero) if (x != y) !ZF
jg (jnle) Jump if greater if (x > y), signed !ZF && !(SF ^ OF)
jge (jnl) Jump if greater or equal if (x >= y), signed !(SF ^ OF)
jl (jnge) Jump if less if (x < y), signed SF ^ OF
jle (jng) Jump if less or equal if (x <= y), signed (SF ^ OF) || ZF
ja (jnbe) Jump if above if (x > y), unsigned !CF && !ZF
jae (jnb) Jump if above or equal if (x >= y), unsigned !CF
jb (jnae) Jump if below if (x < y), unsigned CF
jbe (jna) Jump if below or equal if (x <= y), unsigned CF || ZF
js Jump if sign bit if (x < 0), signed SF
jns Jump if not sign bit if (x >= 0), signed !SF
jc Jump if carry bit N/A CF
jnc Jump if not carry bit N/A !CF
jo Jump if overflow bit N/A OF
jno Jump if not overflow bit N/A !OF

The test and cmp instructions are frequently seen before a conditional branch. These operations perform arithmetic but throw away the result, except for condition codes. test performs binary-and, cmp performs subtraction.

cmp is hard to grasp: remember that subq %rax, %rbx performs %rbx := %rbx - %rax—the source/destination operand is on the left. So cmpq %rax, %rbx evaluates %rbx - %rax. The sequence cmpq %rax, %rbx; jg L will jump to label L if and only if %rbx is greater than %rax (signed).

The weird-looking instruction testq %rax, %rax, or more generally testq REG, SAMEREG, is used to load the condition flags appropriately for a single register. For example, the bitwise-and of %rax and %rax is zero if and only if %rax is zero, so testq %rax, %rax; je L jumps to L if and only if %rax is zero.

You will occasionally see instructions named setFLAG that load the binary value of a condition (0 or 1) into an 8-bit register. For example, setz %al sets %al to 1 if the zero flag ZF is on, and 0 otherwise. There are set constructions for all conditions. set instructions are often followed by zero-extending instructions: setz %al; movzbl %al, %eax sets all of %rax to 1 if the zero flag ZF is on, and 0 otherwise.

Data-movement and control-flow instructions do not modify flags. Oddly, for example, lea does not modify flags (it counts as data movement), though add does (it counts as arithmetic).

Sidebar: C++ data structures

C++ compilers and data structure implementations have been designed to avoid the so-called abstraction penalty, which is when convenient data structures compile to more and more-expensive instructions than simple, raw memory accesses. When this works, it works quite well; for example, this:

long f(std::vector<int>& v) {
    long sum = 0;
    for (auto& i : v) {
        sum += i;
    return sum;

compiles to this, a very tight loop similar to the C version:

        movq    (%rdi), %rax
        movq    8(%rdi), %rcx
        cmpq    %rcx, %rax
        je      .L4
        movq    %rax, %rdx
        addq    $4, %rax
        subq    %rax, %rcx
        andq    $-4, %rcx
        addq    %rax, %rcx
        movl    $0, %eax
        movslq  (%rdx), %rsi
        addq    %rsi, %rax
        addq    $4, %rdx
        cmpq    %rcx, %rdx
        jne     .L3
        rep ret
        movl    $0, %eax

We can also use this output to infer some aspects of std::vector’s implementation. It looks like:

Compiler optimizations

Argument elision

A compiler may decide to elide (or remove) certain operations setting up function call arguments, if it can decide that the registers containing these arguments will hold the correct value before the function call takes place. Let's see an example of a function disassembled function f in f31.s:

subq    $8, %rsp
call    _Z1gi@PLT
addq    $8, %rsp
addl    $1, %eax

This function calls another function g, adds 1 to g's return value, and returns that value.

It is possible that the function has the following definition in C++:

int f() {
    return 1 + g();

However, the actual definition of f in is:

int f(x) {
    return 1 + g(x);

The compiler realizes that the argument to function g, which is passed via register %rdi, already has the right value when g is called, so it doesn't bother doing anything about it. This is one example of numerous optimizations a compiler can perform to reduce the size of generated code.


A compiler may also copy the body of function to its call site, instead of doing an explicit function call, when it decides that the overhead of performing a function call outweights the overhead of doing this copy. For example, if we have a function g defined as g(x) = 2 + x, and f is defined as f(x) = 1 + g(x), then the compiler may actually generate f(x) as simply 3 + x, without inserting any call instructions. In assembly terms, function g will look like

leal    2(%rdi), %eax

and f will simply be

leal    3(%rdi), %eax

Tail call elimination

Let's look at another example in f32.s:

addl    $1, %edi
jmp _Z1gi@PLT

This function doesn't even contain a ret instruction! What is going on? Let's take a look at the actual definition of f, in

int f(int x) {
    return g(x + 1);

Note that the call to function g is the last operation in function f, and the return value of f is just the return value of the invocation of g. In this case the compiler can perform a tail call elimination: instead of calling g explicitly, it can simply jump to g and have g return to the same address that f would have returned to.

A tail call elimination may occur if a function (caller) ends with another function call (callee) and performs no cleanup once the callee returns. In this case the caller and simply jump to the callee, instead of doing an explicit call.

Loop unrolling

Before we jump into loop unrolling, let's take a small excursion into an aspect of calling conventions called caller/callee-saved registers. This will help us under the sample program in f33.s better.

Calling conventions: caller/callee-saved registers

Let's look at the function definition in f33.s:

    pushq   %r12
    pushq   %rbp
    pushq   %rbx
    testl   %edi, %edi
    je  .L4
    movl    %edi, %r12d
    movl    $0, %ebx
    movl    $0, %ebp
    movl    %ebx, %edi
    call    _Z1gj@PLT
    addl    %eax, %ebp
    addl    $1, %ebx
    cmpl    %ebx, %r12d
    jne .L3
    movl    %ebp, %eax
    popq    %rbx
    popq    %rbp
    popq    %r12
    movl    %edi, %ebp
    jmp .L1

From the assembly we can tell that the backwards jump to .L3 is likely a loop. The loop index is in %ebx and the loop bound is in %r12d. Note that upon entry to the function we first moved the value %rdi to %r12d. This is necessary because in the loop f calls g, and %rdi is used to pass arguments to g, so we must move its value to a different register to used it as the loop bound (this case %r12). But there is more to this: the compiler also needs to ensure that this register's value is preserved across function calls. Calling conventions dictate that certain registers always exhibit this property, and they are called callee-saved registers. If a register is callee-saved, then the caller doesn't have to save its value before entering a function call.

We note that upon entry to the function, f saved a bunch of registers by pushing them to the stack: %r12, %rbp, %rbx. It is because all these registers are callee-saved registers, and f uses them during the function call. In general, the following registers in x86_64 are callee-saved:
%rbx, %r12-%r15, %rbp, %rsp (%rip)

All the other registers are caller-saved, which means the callee doesn't have to preserve their values. If the caller wants to reuse values in these registers across function calls, it will have to explicitly save and restore these registers. In general, the following registers in x86_64 are caller-saved:
%rax, %rcx, %rdx, %r8-%r11

Now let's get back to loop unrolling. Let us a look at the program in f34.s:

    testl   %edi, %edi
    je  .L7
    leal    -1(%rdi), %eax
    cmpl    $7, %eax
    jbe .L8
    pxor    %xmm0, %xmm0
    movl    %edi, %edx
    xorl    %eax, %eax
    movdqa  .LC0(%rip), %xmm1
    shrl    $2, %edx
    movdqa  .LC1(%rip), %xmm2
    addl    $1, %eax
    paddd   %xmm1, %xmm0
    paddd   %xmm2, %xmm1
    cmpl    %edx, %eax
    jb  .L5
    movdqa  %xmm0, %xmm1
    movl    %edi, %edx
    andl    $-4, %edx
    psrldq  $8, %xmm1
    paddd   %xmm1, %xmm0
    movdqa  %xmm0, %xmm1
    cmpl    %edx, %edi
    psrldq  $4, %xmm1
    paddd   %xmm1, %xmm0
    movd    %xmm0, %eax
    je  .L10
    leal    1(%rdx), %ecx
    addl    %edx, %eax
    cmpl    %ecx, %edi
    je  .L1
    addl    %ecx, %eax
    leal    2(%rdx), %ecx
    cmpl    %ecx, %edi
    je  .L1
    addl    %ecx, %eax
    leal    3(%rdx), %ecx
    cmpl    %ecx, %edi
    je  .L1
    addl    %ecx, %eax
    leal    4(%rdx), %ecx
    cmpl    %ecx, %edi
    je  .L1
    addl    %ecx, %eax
    leal    5(%rdx), %ecx
    cmpl    %ecx, %edi
    je  .L1
    addl    %ecx, %eax
    leal    6(%rdx), %ecx
    cmpl    %ecx, %edi
    je  .L1
    addl    %ecx, %eax
    addl    $7, %edx
    leal    (%rax,%rdx), %ecx
    cmpl    %edx, %edi
    cmovne  %ecx, %eax
    xorl    %eax, %eax
    rep ret
    rep ret
    xorl    %edx, %edx
    xorl    %eax, %eax
    jmp .L3

Wow this looks long and repetitive! Especially the section under label .L3! If we take a look at the original function definition in, we will find that it's almost the same as, except that in we know the definition of g as well. With knowledge of what g does the compiler's optimizer decides that unrolling the loop into 7-increment batches results in faster code.

Code like this can become difficult to understand, especially when the compiler begins to use more advanced registers reserved for vector operations. We can fine-tune the optimizer to disable certain optimizations. For example, we can use the -mno-sse -fno-unroll-loops compiler options to disable the use of SSE registers and loop unrolling. The resulting code, in f35.s, for the same function definitions in, becomes much easier to understand:

    testl   %edi, %edi
    je  .L4
    xorl    %edx, %edx
    xorl    %eax, %eax
    addl    %edx, %eax
    addl    $1, %edx
    cmpl    %edx, %edi
    jne .L3
    rep ret
    xorl    %eax, %eax

Note that the compiler still performed inlining to eliminate function g.

Optimizing recursive functions

Let's look at the following recursive function in

int f(int x) {
    if (x > 0) {
        return x * f(x - 1);
    } else {
        return 0;

At the first glance it may seem that the function returns factorial of x. But it actually returns 0. Despite it doing a series of multiplications, in the end it always multiplies the whole result with 0, which produces 0.

When we compile this function to assembly without much optimization, we see the expensive computation occurring:

    movl    $0, %eax
    testl   %edi, %edi
    jg  .L8
    rep ret
    pushq   %rbx
    movl    %edi, %ebx
    leal    -1(%rdi), %edi
    call    _Z1fi
    imull   %ebx, %eax
    popq    %rbx

In there is an actual factorial function:

int f(int x) {
    if (x > 0) {
        return x * f(x - 1);
    } else {
        return 1;

If we compile this function using level-2 optimization (-O2), we get the following assembly:

    testl   %edi, %edi
    movl    $1, %eax
    jle .L4
    imull   %edi, %eax
    subl    $1, %edi
    jne .L3
    rep ret
    rep ret

There is no call instructions again! The compiler has transformed the recursive function into a loop.

If we revisit our "fake" factorial function that always returns 0, and compile it with -O2, we see yet more evidence of compiler's deep understanding of our program:

xorl    %eax, %eax

Optimizing arithmetic operations

The assembly code in f39.s looks like this:

leal    (%rdi,%rdi,2), %eax
leal    (%rdi,%rax,4), %eax

It looks like some rather complex address computations! The first leal instruction basically loads %eax with value 3*%rdi (or %rdi + 2*%rdi). The second leal multiplies the previous result by another 4, and adds another %rdi to it. So what it actually does is 3*%rdi*4 + %rdi, or simply 13*%rdi. This is also revealed in the function name in f39.s.

The compiler choose to use leal instructions instead of an explicit multiply because the two leal instructions actually take less space.