SchOS/docs/architecture/kernel.md

SchIFT — the kernel

Long mode, 4-level paging, preemptive multitasking, and more than one processor. What follows is the parts of the kernel that are not a filesystem, a driver or a protocol stack: memory, processes, scheduling, traps and signals.

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Address space

One page table per process; the kernel is mapped into all of them at KERNBASE = 0xFFFF800000000000, so a trap does not switch CR3.

kvmalloc() maps physical 0–4 GB as four 1 GB pages, with the top one marked uncached — that is where PCI BARs and the framebuffer live. This is why 3 GB is the most memory the kernel can treat as RAM, and it is a deliberate simplicity: a direct map with a fixed offset makes P2V/V2P arithmetic rather than a lookup.

kalloc() hands out one 4 KB page at a time from a free list. There is no slab or pool allocator, and the consequence is visible throughout the tree: every subsystem that needs many small objects has invented its own fixed-size table — NPROC processes, NSOCK sockets, NDEVICE devices, NSHM segments. That is why the limits are compile-time constants rather than a design decision about each one.

The free list carries a canary, checked at the top of both kalloc() and kfree(), so a wild write into it panics at the point it is touched rather than handing back a bad pointer that fails somewhere else entirely.

DMA

kernel/core/dma.c owns an arena carved off the top of physical memory before kalloc is given the rest. Devices need physically contiguous memory; kalloc cannot promise it, so the region has to be reserved before there is any competition for it.

Per-CPU state

The per-CPU block's base is in both FS_BASE and GS_BASE, set by seginit():

  +0   self-pointer     required by the x86-64 TLS ABI (%fs:0)
  +8   syscall scratch  trapasm.S parks %rsp here across the stack switch
  +16  struct cpu *
  +24  struct proc *

These offsets are shared with assembly and must not be reordered. Offset 8 is not free: syscall_entry uses %fs:(8), and because FS and GS base point at the same address, that is the same memory as %gs:8.

Reached through mycpu() and myproc(), which compile to a single movq %gs:16 / %gs:24 with no cached base.

__thread is wrong for per-CPU kernel data, and this is the single most expensive lesson in the tree. __thread means per thread, and the compiler is entitled to read the thread pointer once and keep it in a register across calls. But %fs here is per CPU, and a kernel thread changes CPU at every sched(). So the assumption is false for exactly the variables that most need to be right, and the failure is a null dereference deep in the scheduler on any machine with more than one core. include/proc.h carries the disassembly that proves it.

Processes

struct proc in include/proc.h. A fixed table of NPROC entries, allocated by linear scan.

The field order at the top of the struct is fixed by assembly. trapasm.S hard-codes fs_base at offset 0x18, and there is a _Static_assert in proc.c that fails the build if it moves. Anything new goes at the end.

Each process carries: page directory and size, kernel stack, open file table, cwd inode and cwd path string, signal handlers and masks, VMA table for mmap, POSIX identity (pgid, sid, and all six of real/effective/saved uid and gid), and the Linux-ABI syscall map.

Slot reuse

The table recycles slots, and a recycled slot that inherits the previous occupant's state is a whole class of bug that has bitten here twice. allocproc() therefore explicitly clears the signal fields, the exit status, the kill signal, and the identity fields. The failure mode when it did not: deliver_signals() jumping to a totally unrelated earlier process's handler address on this process's very first syscall return.

Kernel threads

kthread_create(fn, name) makes a process that runs fn in kernel mode and never returns to userland. It gets a page directory with only the kernel mapped, no files, no cwd, no parent — it is not in the process tree, cannot be waited for, and must never exit.

The point is that it is a process and not a timer callback: it shares the scheduler's ordinary rotation, so it can hold sleeplocks and block on the disk exactly as a syscall would. The network thread is the current user.

Scheduling

Round-robin over the process table, one scheduler per CPU, each with its own context. swtch() saves callee-saved registers and switches stacks; sched() returns to the scheduler and asserts four invariants on the way — holding ptable.lock, ncli == 1, not RUNNING, interrupts off.

When a full pass finds nothing runnable the CPU halts rather than spinning.

Locking is a single ptable.lock over the whole table. Not per-process: a scheduling operation touches two processes (the one leaving and the one arriving) and two-object locking needs an ordering rule that every path obeys. The same reasoning produced a single lock over the socket table.

sleep / wakeup

xv6's discipline, unchanged because it is correct: sleep(chan, lk) atomically releases lk and blocks; wakeup(chan) makes every sleeper on that channel runnable. The atomicity is what matters — a sleeper that tests its condition under one lock while the waker sets it under another has no interlock and sleeps through the wakeup meant for it.

Traps and syscalls

Interrupts and exceptions enter through vectors.S → trapasm.S → trap(). Syscalls enter through SYSCALL/SYSRET and syscall_entry, which is a separate and faster path.

Every entry from user mode checks a canary at the bottom of the kernel stack page. A deep or nested trap that overflows the 4 KB stack clobbers that word first, so the failure is reported as kstack overflow at the point it happened rather than as arbitrary corruption later.

One configuration detail with an expensive history: MSR_STAR's SYSRET selector field must hold the compat 32-bit code selector, because SYSRETQ computes the returned cs/ss as STAR[63:48]+16/+8. Putting the real 64-bit user selector there makes every sysretq load an invalid pair — which QEMU's TCG does not validate, so it stayed hidden until a hardware interrupt forced a strictly-validated IRETQ.

Signals

Real delivery, not stubs. sys_rt_sigaction() records the handler; deliver_signals() runs in the tail of syscall(), just before returning to user mode, and if a pending unblocked signal has a handler it rewrites the trapframe to enter the handler instead of resuming — stashing the original in proc->saved_sigtf. sys_rt_sigreturn() puts it back.

The return address is not sa_restorer. SA_RESTORER is not defined in this musl build, so sigaction() never fills one in, and pushing 0 as a return address crashes with rip = 0 the moment the handler returns. Instead a 7-byte trampoline (mov eax,15; syscall) is injected onto the user's own stack below rsp, clear of the SysV red zone — the standard technique when the C library supplies no restorer.

Signals without a handler fall back to a default action: termination for the fatal set, dropped otherwise.

PID 1 does not die from a signal it never asked to handle

init_ignores() discards any signal whose disposition is the default action when the target is PID 1 — including SIGKILL. This is what POSIX requires and what Linux does in sig_task_ignored(), and without it the system has a single keystroke that kills it: Ctrl-C signals the foreground process group, init is pgid 1, and any moment where the terminal's foreground group is still 1 lands SIGINT on init. It has no handler, so exit() runs, and exit() panics outright on initproc.

Spamming Ctrl-C at a frozen program is exactly the input that finds that window.

Process groups and sessions

setsid, setpgid, getpgid, getsid, and signal delivery to a whole group. Job control is defined entirely in terms of these — a shell puts each pipeline in its own process group and signals the group, never individual pids — and mksh's job control could not make progress until they existed.

What userspace gets

fork, exec, wait/waitpid/wait4 with WNOHANG and real pid selection, pipes, signals, mmap (anonymous and file-backed), POSIX shared memory, Unix domain sockets, poll/select/epoll, eventfd, timerfd, signalfd.

mmap uses a bump allocator (proc->mmap_next, starting at 0x40000000) so every mapping gets its own disjoint, exactly-sized range. It previously assigned a fixed 8 MB window per VMA slot regardless of the requested length, so a single allocation over 8 MB mapped straight into a neighbouring live mapping.

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What it does not do

• No slab allocator. One page at a time, and every subsystem's fixed-size table is a consequence. This is the single change that would most improve the rest of the kernel.
• No demand paging and no page cache. mmap of a file reads the whole thing in at map time. Writable MAP_SHARED of a file is therefore refused outright rather than faked — see ../plan/foundations.md.
• No dynamic linking. exec handles ET_EXEC only: no PT_INTERP, no ET_DYN, no shared libraries, no dlopen.
• No threads. clone exists but does not share an address space; gettid returns the pid.
• Permissions are tracked but not enforced. uid/gid are maintained honestly so getuid()/setuid() stop lying, and nothing consults them on access.
• No swap, and no memory pressure handling of any kind. kalloc failing is reported up as -ENOMEM and that is the whole policy.
• Nested signal delivery is not supported. in_signal guards against re-entering a handler while one is running.
• Scheduling is round-robin with no priorities, no fairness accounting, and no CPU affinity despite sched_setaffinity existing.