# Scheduling

The scheduler turns danos from a linear "boot then halt" kernel into a **running
multitasking system**. It's **fixed-priority preemptive**: the highest-priority
ready task always runs, and tasks at the same priority take turns. That model is
chosen for [real-time](vision.md) — it's predictable (you can reason about which
task runs when) and its decisions are O(1), unlike a fair-share scheduler.

The scheduler proper (`system/kernel/sched.zig`) is generic; the context switch and new-task
stack setup are architecture-specific (`system/kernel/architecture/x86_64/`, see [arch](arch.md)).

## Tasks

A **task** is a kernel thread: ring-0 code with its own 16 KiB stack (allocated
from the [heap](heap.md)). A task struct holds its saved stack pointer, priority,
state, and a ready-queue link. The currently-running kernel context (kmain)
registers itself as task 0, so there's always something to switch *from*.

## The context switch

Switching tasks means swapping stacks. `switch_context(old, new)` (in `isr.s`)
saves the **callee-saved** registers on the current stack, stores the stack pointer
into the old task, loads the new task's stack pointer, restores *its* callee-saved
registers, and `ret`s — landing wherever the new task was last suspended. Only
callee-saved registers are handled explicitly: to the compiler this looks like a
normal function call, so it already preserves the caller-saved ones itself (this is
the [SysV](sysv.md) convention doing the work).

A **freshly spawned** task has never run, so there's nothing to restore. Its stack
is faked to look as if it had just called `switch_context`: `init_task_stack` lays
down a return address pointing at `task_trampoline` and zeroed callee-saved slots
(smuggling the entry function in via the `r15` slot). When first switched to, the
`ret` lands in the trampoline, which enables interrupts and calls the entry.

## Two ways to switch, one flag discipline

`schedule()` — pick the best task and switch — runs from two places:

- **`yield()`** — a task voluntarily gives up the CPU.
- **`tick()`** — the 1000 Hz [timer](device-interrupts.md) preempts the running
  task. This is what lets a task that never yields still share the CPU.

The subtlety in mixing them is the **interrupt flag (IF)**. The rule: `switch_context`
is always entered with interrupts *disabled* — naturally so inside the timer ISR,
and explicitly (`cli`) in `yield`. Then every task ends up with interrupts enabled
again through whichever path resumes it:

- a task suspended in `yield` re-enables them (`sti`) right after `schedule` returns;
- a task suspended mid-ISR resumes through the interrupt return (`iretq`), which
  restores the `RFLAGS` it had when it was preempted (IF set);
- a brand-new task enables them in the trampoline.

One related detail: the timer interrupt is **acknowledged (EOI) before** its handler
runs, so a handler that switches tasks and doesn't return promptly can't stall the
LAPIC from delivering the next tick.

## Priority selection, in O(1)

Ready tasks live in a **FIFO queue per priority level** (8 levels), plus a
**bitmap** with one bit per non-empty level. Picking the next task is: find the
highest set bit (one instruction), take the front of that level's queue. No list
walking, no scanning — the decision cost is constant regardless of how many tasks
exist, which is what a real-time scheduler needs.

- **Highest priority wins.** A ready high-priority task always runs before a
  lower-priority one.
- **Round-robin within a level.** When a task is descheduled it goes to the *back*
  of its level's queue, so equal-priority tasks share the CPU fairly.

## Affinity: pinning a task to a core

By default a task runs on **any** core — the ready queue above is global, and any
idle core pulls the highest-priority task from it (work-conserving; see
[smp.md](smp.md)). A task can instead be **pinned** to one core with
`spawnOn(entry, priority, cpu)`, giving it an *affinity*: it will only ever run
there, never migrating.

Mechanically, each core has its **own** pinned queue (same 8-level FIFO + bitmap)
alongside the global one. A pinned task is enqueued only into its core's pinned
queue; selection compares the top of the global queue and the running core's pinned
queue and takes the higher priority (still O(1) — two bit-scans and a compare), with
a pinned task winning an equal-priority tie so it can't be starved by global work.
Because every queue is mutated under the [big kernel lock](smp.md), one core enqueuing
into another core's pinned queue is safe.

This is the *explicit-affinity* model (no surprise migration mid-deadline), which is
the more real-time-predictable direction. `spawnOn` refuses to pin to an offline or
out-of-range core — it creates the task unpinned instead, so it still runs somewhere
rather than stranding in a queue no core services, and returns whether the pin took.

## Sleeping and the idle task

A task can **block** — give up the CPU until an event, rather than busy-wait
(busy-waiting is the enemy of a real-time system: it wastes cycles a
higher-priority task should get). The first form is time-based: **`sleep(ms)`**
marks the task blocked with a wake deadline and switches away. On every tick the
timer wakes any task whose deadline has passed (a bounded scan, so it stays
deterministic), which makes it ready again; the scheduler then runs it when its
priority comes up. `sleep` measures its deadline on the [calibrated
clock](device-interrupts.md), so it's real time.

When *every* task is blocked, something still has to run — so there's an **idle
task** at the lowest priority that just `hlt`s until the next interrupt (see
[halting.md](halting.md)). Because it's always runnable, the scheduler always has a
task to pick, and the "nothing to run" case never arises.

## Event-based blocking

The other form of blocking is waiting for an **event** rather than a duration. A
**wait queue** is a set of tasks parked until something happens: `wait(wq)` blocks
the caller on it, `wake(wq)` moves the highest-priority waiter back to ready
(preempting if it now outranks the running task). A task links into a wait queue
through the same field the ready queues use — it's in exactly one queue at a time.
These are the primitives locks, semaphores and [IPC](ipc.md) are built on.

Blocking safely needs **composable critical sections**. A blanket `cli`/`sti` pair
doesn't nest: an IPC channel that `cli`s and then calls `wait` would have `wait`'s
`sti` re-enable interrupts too early, mid-operation. So the blocking primitives use
`saveInterrupts` / `restoreInterrupts` — capture the interrupt flag, disable, and
later restore *only if it was set* — which nests correctly. The invariant that
makes it all work: `schedule()` is always entered with interrupts disabled, so a
task always resumes from a switch with interrupts disabled and can restore its
caller's state.

## Verifying it

Three tests (see [testing.md](testing.md)) prove the guarantees:

- **`sched`** spawns three tasks that busy-loop *without ever yielding*. They all
  make progress — which can only happen if the timer is **preempting** between them
  and the context switch is correct (nothing yields voluntarily).
- **`priority`** (with preemption off, for determinism) spawns tasks at three
  priorities; they run and exit **highest-priority first** — `[6, 4, 2]`.
- **`sleep`** blocks a task for 50 ms and checks the elapsed time on the clock — a
  real block (the idle task runs meanwhile), not a busy-wait.
- **`event`** blocks a task on a wait queue; waking it (from another task) resumes
  it, and since it's higher priority it preempts immediately.

## What's next (not done here)

- **Priority inheritance.** Once tasks block on shared resources (locks, IPC),
  danos will need it to bound priority inversion — a [real-time](vision.md)
  requirement.
- **Task exit / a reaper.** `exit` currently leaks the task's stack; nothing frees
  finished tasks' memory yet.
- **Per-address-space tasks.** Today all tasks share the kernel address space. User
  processes will each get their own, switching page tables (CR3) on the context
  switch.
