Two real bugs visible in a normal `zig build run-x86-64` boot (but not in the
display-demo test, which spawns no input service):
- Two cursors. The demo drew its own cursor layer while the display service
now draws one too (its mouse-listener thread). The demo's went through the
client IPC protocol and lagged; the service's is in-process and tracks
tightly — "one responds better than the other."
- Animation frozen until the mouse moves. The demo called a *blocking*
`mouse.next()` (ipc_reply_wait) inside its animation loop, so the sliding
box advanced only one frame per mouse event. In the display-demo test
there is no input service, so subscribeMouse() returned null and the loop
ran free on its 30ms timer — which is exactly why the test passed while
the real boot was broken.
The compositor now owns the cursor (docs/display.md), so the demo should draw
none and read no input: it becomes a pure client-animation proof whose loop is
independent of the mouse. Removes its cursor layer, mouse subscription, and the
blocking read.
Also harden displayDemoTest to spawn the `input` service alongside the demo
(matching real boot): a client that blocks its animation loop on a mouse read
would now stall before `display-demo: ok` and fail the test, instead of
passing because no input service happened to be present.
Verified: display / display-service / display-demo / display-cursor all green.
The display's first use of threads (docs/threading.md, docs/display.md). The
compositor stays the single owner of the framebuffer — only the main
service.run loop touches the backend and layer stack — and a dedicated
mouse-listener thread runs beside it:
- Listener: blocks on input.subscribeMouse(), accumulates relative dx/dy
into an absolute cursor position clamped to the screen, and hands it to
the compositor. It never touches the compositor, so no lock guards the
framebuffer; a parked next() lets the core halt.
- CursorChannel: a single-slot latest-value cell under a Thread.Mutex (the
renderer wants where the cursor is now, not a replay of deltas), with a
coalesced self-ipc.send poke that wakes the main loop — parked in
replyWait — as a message-notification. At most one poke is queued while
the last is undrained, so a fast mouse can't flood the endpoint.
- Render: the cursor is a top-z compositor layer; on the poke the main loop
moves it via configure + present (which damages old + new footprints).
The display binary opts into threads (addThreadedUserBinary), and
input-source gains a "mouse" mode that publishes pure motion to drive it.
Two kernel-level findings this surfaced, both fixed:
1. IPC handles do not cross threads. The handle table lives on the Task, so
the listener can't reuse the main loop's endpoint handle — it
ipc.lookup(.display)s its own handle to the same endpoint to poke through.
2. Concurrent IPC from two threads raced unlocked kernel state. The display
is the first process issuing IPC syscalls from two threads at once, which
exposed a data race (flaky #GP in installEntry): create_ipc_endpoint /
ipc_register / ipc_lookup allocate from the kernel heap and mutate the
global registry, endpoint refcounts, and handle tables without the big
kernel lock. They were safe only while a process couldn't race itself.
They now sync.enter() like call/reply_wait/send already did (the kernel
heap has no lock of its own yet — heap.zig: "a lock comes with
threads/SMP" — so the big lock keeps its callers serialized).
Test: -Dtest-case=display-cursor (smp:4) spawns the input service, the
threaded display, and input-source in mouse mode; asserts the display's
"cursor tracking mouse ok" marker once the cursor has tracked a run of motion
end to end. Verified green 6/6 under stress (the race hit ~1-in-4 before the
lock fix) and in the full 29-case QEMU guardrail suite; zig build test clean.
The "fs.base" spelling read like a field/submodule access, but no such
identifier exists — it meant the x86_64 FS segment base (the IA32_FS_BASE
MSR). Two problems fixed:
- Arch-neutrality: in generic docs, the runtime, and the plan, the mechanism
is now named by its arch-neutral concept — the "thread pointer" — matching
the already-renamed `thread_pointer` Task field, `set_thread_pointer`
syscall, and `architecture.setThreadPointer` fn. x86-specific spots keep
the precise names: `IA32_FS_BASE` (the MSR), `%fs:8`/`%fs:0`, variant-II.
- Stale identifiers: the M10 section in threading-plan.md still referenced
`fs_base` on Task and `architecture.setFsBase` — both renamed away in the
arch-neutral pass. Corrected to `thread_pointer` / `setThreadPointer`.
The x86-only `thread-tls` test (which really does write `%fs:8`) now says
"FS base" (no dot) consistently, matching the established form already in
tests.zig. The matched serial markers ("thread-tls: ok" / "thread-tls:
FAIL") are unchanged; only a non-load-bearing FAIL parenthetical was
reworded.
Verified: zig build clean, thread-tls passes.
Expand the abbreviations flagged in docs/coding-standards.md (names spelled
out in full unless an acronym) across the kernel, runtime, ABI, tests, and
docs:
aspace -> address_space (AspaceRef -> AddressSpaceRef, retainAspace ->
retainAddressSpace, loaded_aspace -> loaded_address_space, the
liveAspaceCount/aspaceDestroyCount test hooks, etc.)
vaddr -> virtual_address
paddr -> physical_address
The kernel test case and its serial markers are renamed to match:
aspace-refcount -> address-space-refcount (kernel dispatch string and
test/qemu_test.py case name kept in sync). Prose in docs uses the natural
"address space"/"virtual address"; backticked field/identifier references
use the code spelling.
Also expand the bare "AS" abbreviation in three ABI comments and reframe the
set_thread_pointer ABI/handler docs to lead with the arch-neutral concept
(user-space TLS thread pointer; x86_64 IA32_FS_BASE, aarch64 TPIDR_EL0)
rather than x86 FS-first, matching scheduler.zig's existing framing.
Foreign ABI names preserved: the ELF p_vaddr field and mmap/mmio remain.
Verified: zig build, zig build test, and the full 25-case QEMU guardrail
suite all green.
The M10 TLS work leaked x86 naming into the generic kernel: Task.fs_base,
PerCpu.loaded_fs_base, and architecture.setFsBase. The *mechanism* was already
abstracted (the generic scheduler calls through the architecture layer; the
wrmsr IA32_FS_BASE lives in architecture/x86_64/cpu.zig), but the *names* would
force an aarch64 port to implement a 'setFsBase' that writes TPIDR_EL0.
Rename to the neutral concept: Task.thread_pointer, PerCpu.loaded_thread_pointer,
architecture.setThreadPointer (x86_64 impl writes IA32_FS_BASE; aarch64 -> TPIDR_EL0).
Also neutralise the user_arg comment (first argument register, rdi on x86_64).
No behaviour change; thread-tls/thread-mutex/smp/process-kill + host tests pass.
runtime.Thread.RwLock (reader-preferring, lock/tryLock/unlock +
lockShared/tryLockShared/unlockShared) and WaitGroup (start/finish/wait), both on
the existing Mutex/Condition.
A compile-time Futex seam gated on builtin.os.tag: the futex syscalls on danos, a
spin+yield mock off-target (Zig 0.16 has no std.Thread.Futex; wake is a no-op
since the state machines re-check). thread.zig is wired into zig build test, so
Mutex/RwLock/WaitGroup run as host unit tests with real std.Thread threads (test
blocks compile only under test, so std.Thread there is fine on freestanding).
thread-rwlock QEMU case: 2 writers set both halves of a value under the exclusive
lock while 3 readers check they match under the shared lock; zero half-write
observations across ~150k reads.
Marks Phase 2 (M7-M11) built. threading.md/threading-plan.md status updated.
Gate: host zig build test covers the sync primitives; thread-rwlock PASS (3x);
full Done gate 26/26 (whole thread-* suite + guardrail); build clean.
Each thread gets its own x86_64 thread pointer (FS base) for user-space TLS.
Task.fs_base is restored on every context switch only when it changes (same
conditional-load discipline as CR3; architecture.setFsBase -> wrmsr
IA32_FS_BASE). New set_thread_pointer=44 syscall sets the caller's fs_base and
loads it now. The kernel never touches FS, so no swapgs complication.
The runtime lays a small per-thread TLS block at the top of each thread's stack
(self-pointer at %fs:0 + scratch) and the thread trampoline calls
set_thread_pointer before any user code — so every spawned thread has a private,
switch-stable thread pointer, reclaimed with the stack.
thread-test tls mode: two threads write unique markers to their own %fs:8 and,
after both wrote, read back — a shared fs.base would clobber one (cross-talk).
Deferred: the Zig threadlocal *compiler* layer (ELF variant-II PT_TLS + linker
sections + template copy) — high-uncertainty, no consumer today; this lands the
load-bearing per-thread fs.base it builds on. See docs/threading-plan.md M10.
Gate thread-tls PASS (3x); full guardrail 25/25; build + host tests clean.
join no longer needs a per-thread IPC endpoint. New thread_join(tid) syscall
blocks the caller until the task with id tid exits; the exit paths call
wakeJoinersLocked. join only reclaims the joined thread's USER stack, which the
thread vacates the instant it enters the kernel to exit, so waking at exit time
(not reap time) is safe — no reaper/aspace juggling or user-memory write, and
equally std-shaped (like pthread_join). thread_spawn drops the exit-endpoint arg
(runtime passes no_cap). thread-test's join mode runs 40 spawn+join cycles that
would exhaust the 16-slot handle table under the old endpoint scheme.
Also harden the M8 reaper: its single per-core reap slot could be overwritten by
a second death on that core before draining (a fresh-task/SMP timing window), an
intermittent one-stack leak that made task-reap ~20% flaky. Replace it with a
per-core reap LIST plus a .reaping task state so a pending slot can't be reused
before its stack is freed. task-reap now 11/11 isolated.
Deferred: detached-thread user-stack reclaim (still at process exit, as in M3).
Gate thread-join PASS (3x); full guardrail 26/26; build + host tests clean.
A dead task's kernel stack was leaked (no reaper), so every process/thread death
bled kernel memory. Now exit()/exitUserLocked record the dying task in a per-core
reap_after_switch slot and switch away; the task that resumes on that core frees
the stack in switchTo's tail (on its own stack, lock still held so the slot can't
be reused). reapKillPendingLocked drains the slot on the timer tick as a safety
net for the fresh-task case (a fresh task enters via the trampoline, bypassing
switchTo's tail). A task killed while not running is freed directly in
destroyTaskLocked. live_stack_bytes is the observable.
Fixed a migration bug this exposed: the post-switchContext reap read the pc
parameter, but a migrated task carries a stale pc in its saved switchTo frame ->
it freed the wrong core's pending stack (a #GP under SMP). Re-fetch thisCpu()
after the switch.
Gate task-reap PASS (5x isolated, 2x in the 24-case batch); full guardrail 24/24
incl. fault-recovery/supervision/process-kill/smp/affinity; build + host green.
Move the mmap/mmio grant-arena cursors off Task into the per-address-space object
(scheduler aspace_refs, exposed via aspaceMmapNextPtr/aspaceDeviceMapNextPtr), so
sibling threads in one address space hand out disjoint grants. systemMmap reserves
a range under a brief lock then maps per page under a short-held lock (not the
whole grant): the big lock runs with interrupts disabled, so pinning it across a
multi-MiB memset+map froze other cores. Guard the runtime heap's rawAlloc/rawFree
with a Thread.Mutex, gated on !single_threaded so ordinary binaries compile it out.
thread-test gains an alloc mode: 4 threads x 500 alloc/fill/verify/free cycles;
any overlap between concurrent allocations is caught by the pattern check.
Also fix the affinity guardrail: its 3-billion-iteration busy-loop had
codegen-dependent wall-time (adding a function to tests.zig swung it ~4s -> ~63s
and timed it out). Reworked to wait on the wall clock instead.
Gate thread-alloc PASS (3x); full guardrail 23/23 green; build + host tests clean.
New thread_self=42 syscall backs runtime.Thread.getCurrentId (the calling
thread's kernel task id). thread-test gains an id mode: two workers read
getCurrentId and the main thread confirms all three ids are non-zero and
distinct. Per-thread threadlocal TLS is deferred by design (no consumer; it would
need context-switched fs.base for an unused feature), as are RwLock/WaitGroup.
Marks the threading feature built (M1-M6): threading.md + docs/README.md status
updated, all thread-* cases wired into qemu_test.py.
Gate thread-id PASS; full suite green: 21/21 (thread-spawn/join/futex/mutex/id,
aspace-refcount, + 15 guardrail cases), zig build clean, zig build test green.
runtime.Thread.Mutex is the classic three-state futex mutex (unlocked/locked/
contended): the fast path is a single CAS and only a contended lock enters the
kernel. Condition is a futex sequence counter (wait/timedWait/signal/broadcast,
spurious wakeups allowed, use in a predicate loop); a signal racing the unlock
bumps the seq so it is never missed. Semaphore is permits guarded by
Mutex+Condition. All mirror std.Thread's shapes, ported onto runtime.Thread.Futex.
thread-test gains a mutex mode: 2 producers + 2 consumers move 2000 unique items
through an 8-slot ring (small enough that both sides block); the consumed
checksum and tally match exactly, proving the lock and condvars correct under
real cross-core contention.
Deferred with rationale (see docs/threading-plan.md): migrating join to a futex
completion word needs kernel clear-on-exit (else use-after-free munmapping a live
stack); host unit tests need a mockable Futex seam.
Gate thread-mutex PASS (3x); 17 guardrail/thread cases green; build + host tests
clean.
New private syscalls futex_wait(addr, expected, timeout_ns)=40 and
futex_wake(addr, count)=41. A waiter is a .blocked task tagged with
Task.futex_addr (no queue linkage); futex_wait reads the user word under the big
lock and parks only if it still equals expected, so a concurrent wake can't slip
between the check and the block. futex_wake scans the task table and readies up
to count waiters in the same address space. A timed wait also sets wake_at so the
existing wakeExpired times it out; futex_addr staying non-zero (only futex_wake
clears it) distinguishes timeout from a real wake. Waiters park in-kernel, so an
idle core still halts (no busy-wait).
runtime.Thread.Futex mirrors std.Thread.Futex (wait/timedWait/wake). thread-test
gains a futex mode: a waiter parks, the main thread wakes it (serial order
waiting/waking/woke, asserted by the case regex), and timedWait reports a
timeout.
Gate thread-futex PASS (3x); 18 guardrail cases green incl. sleep/event/ipc
blocking paths; build + host tests clean.
thread_spawn takes a 4th arg, an exit-endpoint handle: spawnThreadSupervised
resolves and refcounts it under the spawn lock (like spawnProcessSupervised), so
a thread's death posts a child-exit notification carrying its tid. runtime
Thread.join blocks in replyWait on that (private) endpoint for its tid, then
munmaps the stack; detach relinquishes the join (stack reclaimed at process
exit, for now). New current_core=39 syscall + Thread.currentCore() lets a worker
observe which core it ran on.
The closure now lives at the top of the thread's own (private) stack instead of
the heap, so spawn/join never touch the not-yet-thread-safe runtime heap.
thread-test gains a join mode: 4 workers x 100k atomic increments, joined, with
counter == N*K and >1 core stamped (real parallelism), plus a detached worker.
Gate thread-join PASS (4x, non-flaky); 17 guardrail/M1/M2 cases green; build +
host tests clean.
A thread is a task sharing the caller's address space. New private syscalls
thread_spawn(entry, stack_top, arg)=37 and thread_exit=38: thread_spawn goes
through scheduler.spawnThread (retains the shared aspace), thread_exit ends the
task like a process exit(0) (terminateCurrent -> releaseAspace, so the space
survives while siblings hold it). The closure pointer reaches the new thread in
rdi via a new jump_to_user_arg asm path and a per-task user_arg (0 for a normal
process, whose _start ignores it) - so the runtime trampoline is a plain C-ABI
Zig function, no naked asm.
runtime.Thread (library/runtime/thread.zig) mirrors std.Thread.spawn: mmap a
stack, heap-allocate the args closure, hand the kernel the trampoline + closure.
addThreadedUserBinary opts a binary into single_threaded=false; thread-test is
the first, and proves a worker runs in the shared address space via a shared
global the main thread polls.
Gate thread-spawn PASS; 16 guardrail cases green (incl. args/init/process on the
new jump_to_user_arg path) + aspace-refcount; build + host tests clean.
Route address-space lifetime through a refcount keyed by the page-table root
(scheduler.zig aspace_refs): retainAspace on the spawnUserLocked success path,
releaseAspace from both teardown paths (exitUserLocked, destroyTaskLocked),
destroying the space only when the last task on it exits. Behaviour is identical
today (every space has exactly one task); this is the foundation shared-address-
space threads (docs/threading.md) build on.
Test-observable liveAspaceCount/aspaceDestroyCount + a new aspace-refcount kernel
self-test and QEMU case: spawn and reap 5 ring-3 probes, assert live spaces return
to baseline and destructions advance by exactly 5 (destroyed once each, no leak,
no double-free). Gate passes; 13 guardrail cases green; build + host tests clean.
The kernel enabled SSE at boot and both kernel and userspace keep live values in
XMM (LLVM emits movdqu/movaps for >=16-byte struct copies, plus floats and SIMD),
yet the kernel never saved the SSE/FPU register file anywhere — not across
switch_context, not across the syscall boundary, not across interrupts. Any value
the compiler parked in an XMM register across a kernel entry could be silently
clobbered by kernel code, or by whatever the scheduler ran while the task blocked:
a whole-kernel, timing-dependent data-corruption bug.
It was the real root cause of the "device_enumerate corruption" Heisenbug: the
display service read garbage framebuffer geometry (height=0) because findDisplay
held the 16-byte .display field live in xmm0 across the claim/mmio_map syscalls,
and a timer preemption to the busy device-manager clobbered it. The tell that it
was register-only: memory always read correct, and the bug vanished whenever an
added syscall spilled the value to the stack.
Fix: fxsave/fxrstor the register file in the asm stubs (isr_common and
syscall_entry), right after pushing the GP trap frame — before any Zig kernel code
can touch XMM — and right before the pops. rbx bridges the exact rsp across the
call to interruptDispatch: it is callee-saved, so it survives even a blocking
dispatch that context-switches away and back, and `and $-16,%rsp; sub $512,%rsp`
gives fxsave its 16-byte-aligned scratch on the kernel stack. Context-switch-time
save/restore alone is not enough — kernel code between the interrupt and switchTo
already clobbers XMM.
The commit-58927ed Gop.init workaround (copy scalar geometry fields rather than the
whole descriptor by value) is now redundant but harmless; left in place. Deferred:
a fresh task inherits the previous task's XMM (minor info-leak / nondeterminism),
and the fxsave runs on every interrupt including ring0->ring0.
pciScanTest put three [64]DeviceDescriptor arrays on the stack; at 328
bytes each that is a 66,096-byte frame, larger than the 64 KiB bootstrap
stack the boot context runs on, so the prologue overflowed the stack and
triple-faulted on entry before the test could print anything — the CPU
reset with no exception line. Reuse one descriptor buffer (frame ~21 KiB).
With the fault gone, the restart-drill checks proved flaky. They polled
process.write_buffer (which holds only the latest write-syscall message)
for the transient restarting/re-scan lines, which the concurrent
class-driver output overwrote before the starved low-priority poller could
observe them. The kernel broker count can't witness the restart either:
register is idempotent and the table has no unregister, so the count is
invariant across kill/prune/respawn.
Follow the sibling restart drills (usbReportTest/driverRestartTest): assert
the ordered drill in the harness regex over the full serial log (a
backreference requires the respawn to re-scan the same count), and keep only
race-free broker checks in the kernel — empty before the scan, populated
once it settles, never growing past N — sampled via scheduler.sleep at a
fixed cadence instead of a busy-poll.
The compositor now survives the virtio-gpu driver dying and re-attaches when device-manager
restarts it — the last piece of display v2.
Three parts:
- The driver hellos the device manager (role: bus). It never did, so the manager — which
spawns it supervised and expects a hello — was stopping it at the 3s hello deadline every
run (the gate markers just printed first). Now it is properly supervised: not stopped for
silence, and restarted on death.
- A kernel IPC fix so a call to a dead service errors instead of hanging. An endpoint records
its owner; when that task dies, its registered endpoints are marked dead (and any parked
senders woken with -EPEER), so ipc_call returns -EPEER rather than blocking on a reply that
will never come. Without this the compositor's first present after the driver died blocked
forever. General robustness — any client of any service benefits.
- The compositor re-attaches. Its .scanout calls now fail cleanly (caught), freezing the last
frame; when the restarted driver re-announces, attach_scanout detects the backend is already
virtio and logs "scanout re-attached", mapping the fresh shared surface and re-looking-up
.scanout. (The previous shm mapping leaks — no shm_unmap syscall yet — but its frames are the
dead driver's, reclaimed on exit.)
- device-manager gains a "test-scanout-restart" mode (like test-usb-restart) that kills the
virtio-gpu driver once after it hellos; the displayReattachTest kernel scenario drives it.
Gate: python3 test/qemu_test.py display-reattach — "scanout upgraded to virtio-gpu" then
"scanout re-attached", no CPU exception, passing 3/3. host tests, ipc/ipc-call/ipc-cap,
supervision, shm, display-service, display-demo, virtio-gpu, display-native, and
display-modeset all pass; default zig build clean. v2 (V1-V6) complete.
The native backend can now change resolution and presents tear-free.
Mode-setting without churn. The driver sizes its scanout resource + shared surface to the
largest mode it offers and treats a mode change as re-pointing the scanout rectangle within
that surface — so the resource, its backing, and the shared mapping never change, and the
surface's row stride (the max width) is fixed while the active width/height move. The
compositor is handed that stride in the announce and composes at it; a smaller mode just
paints the top-left rectangle. This sidesteps the surface re-share a true resolution change
would otherwise need (the service harness can't reply with a capability).
- scanout-protocol gains get_modes + set_mode; the driver offers {640x480, 800x600} and
re-points set_scanout on set_mode.
- backend.VirtioGpu carries the surface stride, exposes modes()/setMode(), and reports
canModeSet = hasVsync = true.
- runtime.display gains modes()/setMode() (display-protocol get_modes/set_mode, forwarded to
the backend) — the client-facing API.
- EDID: the driver negotiates VIRTIO_GPU_F_EDID when the device offers it, reads the monitor's
EDID, and logs its preferred mode (parsed from the first detailed timing descriptor).
- vsync: every resource_flush is fenced (VIRTIO_GPU_FLAG_FENCE); the device signals the fence
when the frame is on screen, which the used-ring ack the synchronous present already waits
on gates — so a completed present is a tear-free one.
After the native upgrade the compositor runs a one-shot mode-set self-check: query the modes,
switch to a different one, re-composite, and confirm the backend reports the new geometry —
the gate's markers.
Gate: python3 test/qemu_test.py display-modeset (reuses the display-native boot) — "display:
mode set to 800x600, verified" + "display: vsync present ok", passing 3/3. host tests,
display-service, display-demo, shm, virtio-gpu, and display-native still pass.
The compositor now boots on the GOP framebuffer and upgrades to the virtio-gpu driver
the moment it announces itself — the pluggable-scanout payoff.
The shared surface. The scanout resource is an shm region the driver creates
(shm_physical, a new syscall, hands it the guest-physical for attach_backing) and passes
to the compositor as a capability. The compositor maps it and composites straight into
it: on x86 DMA is cache-coherent, so the cacheable shared pages the CPU paints are exactly
what the device transfers-and-flushes — no copy, no explicit flush.
The handshake. After bring-up the driver looks up .display and sends attach_scanout with
the geometry + the surface capability. The compositor maps the surface, looks up the
driver's .scanout endpoint itself (the driver registered it — no need to pass it), switches
to backend.VirtioGpu, and re-composites the current frame. present() over the native
backend is a present request on .scanout -> transfer-to-host + resource flush. The first
native present is deferred to a one-shot timer: presenting inline from the announce handler
would deadlock, since the driver is still blocked on our reply and not yet serving .scanout.
After it lands, the compositor reads a pixel back from the shared surface to confirm the
frame reached the device's backing.
- shm_physical (syscall 36) + runtime.shm.physical.
- scanout-protocol (the compositor->driver present channel), separate from the
client-facing display protocol; the display protocol gains attach_scanout.
- backend.VirtioGpu joins backend.Gop in the tagged union; select() still boots GOP.
- the virtio-gpu driver's scanout backing is now shm (was DMA); it announces + serves
.scanout present requests (transfer-to-host + flush of the shared surface).
Also fixes a latent framebuffer-geometry corruption the display service hit only when it
enumerated the device tree alongside a busy device-manager: Gop.init now captures the
geometry into a small value the instant device_enumerate returns (rather than re-reading
the 328-byte descriptor across the later claim/mmio_map syscalls) and retries on a zero
geometry. The underlying device-table clobber is a separate kernel bug, tracked apart.
Gate: python3 test/qemu_test.py display-native (QEMU -device virtio-gpu-pci) — "display:
scanout upgraded to virtio-gpu" + "display: native present verified" + "display-demo: ok",
passing 3/3. host tests, display-service, display-demo, shm, and virtio-gpu still pass.
The first native scanout backend's driver half. The device manager matches the
virtio-gpu PCI function by its Display/Other class triple and spawns the driver,
which claims the device, confirms vendor 0x1AF4/device 0x1050 from config space,
enables memory-space + bus mastering, and walks the virtio vendor capabilities to
find the common-config and notify structures in its BAR.
From there it is the standard modern-virtio bring-up: reset, negotiate VERSION_1,
stand up the control virtqueue in coherent DMA, then drive the GPU end to end —
RESOURCE_CREATE_2D, ATTACH_BACKING (a coherent DMA buffer; V4 swaps in the shm-
shared surface), SET_SCANOUT, paint a known pattern, TRANSFER_TO_HOST_2D,
RESOURCE_FLUSH, and wait for the device's used-ring ack. Reading the backing back
proves it is CPU-visible RAM; the ack proves the device consumed the frame —
together the automated stand-in for "it's on screen", no screenshot.
- system/drivers/virtio-gpu/: the driver, plus virtio-gpu-protocol.zig (control
commands) and virtio-pci.zig (the 1.0 PCI transport + split-virtqueue), both with
host-tested struct sizes.
- device-manager matches the display/other class triple to "virtio-gpu"; the driver
self-confirms the vendor/device id, since the class alone cannot distinguish it.
- ServiceId.scanout (11): the driver registers it so the compositor finds it in V4.
Gate: python3 test/qemu_test.py virtio-gpu (QEMU -device virtio-gpu-pci) — the
device-manager stack discovers the function, the driver brings up a 640x480 scanout
and flushes a test pattern: "virtio-gpu: scanout 640x480 online" + "flush acked,
pixel check ok". host tests, display-service, shm, and device-list still pass.
Note: the pre-existing pci-scan case triple-faults on main (verified at 88ad432,
before this change); it is unrelated and tracked separately.
Generalize capability passing from endpoints to memory objects. The per-task
handle table now holds kind-tagged entries (scheduler.HandleObject{kind, ptr});
closeHandles and shareCapability dispatch by kind, so a shared-memory object
rides an ipc_call send_cap exactly like an endpoint and is refcount-freed only
when its last capability drops.
- shm_create(len) -> vaddr, handle: contiguous, zeroed, cacheable frames wrapped
in a refcounted ShmObject, mapped into the caller's shm arena (PML4[230]).
- shm_map(cap) -> vaddr: map the same physical pages into a receiver that got the
capability. mapUserSharedInto maps WB-cacheable + device_grant, so a sharer's
teardown never frees the shared frames — the object owns them.
- runtime.shm: create(len) -> Region{ptr, handle, len}, map(handle) -> ptr.
Gate: qemu_test.py shm — shm-client creates a region, writes a pattern, passes
its capability to shm-server, which maps it and reads the same bytes back
(shm: shared 4096 bytes ok). ipc/ipc-call/ipc-cap/supervision/dma/usermem/
display-service and host tests all still pass — the handle change broke no IPC.
Extract scanout from the compositor into backend.zig — a `Backend` tagged union
with info()/surface()/present(damage) and canModeSet/hasVsync flags. The v1 GOP
path becomes `backend.Gop` (claim the display node, WC-map the LFB, keep the
cacheable back buffer, present = the damage-rect WC copy); display.zig now
composes into backend.surface() and calls backend.present(damage), with no LFB or
framebuffer geometry left in the compositor core. The selection decision is the
pure chooseKind(native_available), split from the syscall-bound bring-up, ready
for the native-if-present branch at V4.
Pure refactor: display-service, display-demo, and zig build test all pass
unchanged.
init supervises its boot services (spawned against an exit endpoint) but the
child-exit handler was `if (got.isNotification()) continue;` — it silently
dropped a dead service. Now init restarts it: on a child-exit notification it
finds the service, and unless it exited cleanly (chose to stop) or has hit the
crash-loop cap (maximum_restarts), respawns it and logs the death + reason. The
reincarnation half of resilience (docs/resilience.md) at the service level, the
counterpart to the device manager's driver restarts. A `shutting_down` flag
skips restarts during the orderly stop sequence, whose child deaths are expected.
Two gaps a real-hardware crash exposed, both in onException's terminal path (and
the panic path):
- The report was anonymous. Add the faulting task's id + name and whether it
trapped in ring 3 (a user process) or ring 0 (the trusted base) — so a fatal
fault says WHAT crashed and WHERE, not just the vector. scheduler gains
currentIdSafe/currentNameSafe (early-boot-guarded, like currentCpuIndex, so the
reporter can't fault a second time).
- halt() never released the big kernel lock, so a core that died holding it
deadlocked every other core spinning in acquire() — the whole machine hangs,
not just the one core the design promises. The BKL now records its owner
(architecture.cpuLocal(), a unique per-core token); sync.releaseIfHeldHere()
frees the lock only if this core holds it, called before halt on both fatal
paths. Caveat: if we held it mid-mutation the shared state may be inconsistent,
but letting the other cores + the supervisor keep running is strictly more
recoverable than a guaranteed total hang.
Spawn the display-demo client alongside the compositor so a normal boot shows
the moving scene (wallpaper + sliding rectangle + cursor) — the GUI track's
visible payoff. A demo fixture: drop it from boot_services to boot to a bare
compositor.
Now that the user-space display service owns the framebuffer, the bootstrap
console shrinks to fatal-only. status/statusPrint go to the diagnostic log
alone, so routine boot output no longer scribbles on a screen the compositor is
about to paint — and a driver's recoverable fault report stays off it too. A new
fatal/fatalPrint keeps panics and kernel-mode faults on screen, forcing the
console back on so a dying machine's last words show even over a live display.
console.zig now documents its early-boot + fatal-fallback role.
Drive the compositor from a separate process, proving the pipeline end to end.
- runtime.display: a Layer handle (fill / blitTile / configure / damage /
destroy), createLayer, and a color(r,g,b) helper that caches the mode and
packs via protocol.pack. Coordinates are signed over the u32 wire fields
(@bitCast both ways), so a layer may sit or move partly off-screen.
- system/services/display-demo: the input-source analog for the compositor — a
full-screen wallpaper, a rectangle it slides back and forth (moved by
configure each frame, so the damage-driven present repaints old + new), and a
cursor. Presents in a loop paced by runtime.time. Wired into build + initrd.
Fix this surfaced: protocol.message_maximum was 4096, but the kernel caps every
IPC message at MESSAGE_MAXIMUM = 256, so replyWait rejected the oversized receive
buffer with -E2BIG and the serve loop had been spinning since D2 (invisibly, as
those gates matched init-time heartbeats). Set it to 256; blit_tile is now
explicitly a small-tile path (larger bitmaps are the deferred shm surface).
Gate: `python3 test/qemu_test.py display-demo` — the demo drives frames of
motion through the layer client API and logs `display-demo: ok`. Regression:
zig build test, display (D1), display-service (D2/D3), and default zig build.
Turn the service into a real compositor. A layer is a server-owned surface (its
own mmap'd cacheable buffer) with a screen position, z-order, and visibility.
Clients create layers, draw into them by command, mark damage, and present; the
compositor repaints only the damaged region — clear to the wallpaper, paint the
visible layers bottom-to-top (z-sorted), flush that rectangle back -> front (WC).
- compositor.zig: the pure, host-tested core — Rect (intersect/unite), Surface,
fillRect, composite (opaque, clipped to a damage rect), blitTile (unaligned-
safe read of a client tile). No syscall/runtime dependency.
- protocol.zig: pack(format, r, g, b) — native pixel encoding for rgbx/bgrx, the
shared colour vocabulary of client and server. Host-tested.
- display.zig: the layer table + create/configure/destroy/fill_rect/blit_tile/
damage/present ops wired onto the compositor, plus a damage-accumulating present.
- Startup self-check: two overlapping layers composited on the real framebuffer,
read back to confirm the overlap shows the top layer and outside shows the
bottom — logs `display: compositor self-check ok`.
Gate: `zig build test` green (compositor + pack), and the display-service case's
self-check passes on hardware. Both new pure modules added to the test loop.
Stand up /system/services/display: a ring-3 process that claims the framebuffer
D1 seeded, maps it write-combining as the front buffer, allocates a cacheable
back buffer, and presents composed frames. The GUI track's compositor, reached
by name over ServiceId.display (= 9).
- protocol.zig: the display wire protocol (info/create_layer/configure_layer/
destroy_layer/fill_rect/blit_tile/damage/present). `info` and a whole-screen
`present` are live; the layer ops fail-stub until D3.
- display.zig: enumerate -> claim -> mmio_map(WC) the LFB, mmap a cacheable back
buffer, clear it and present it (proving the double-buffer path), then serve.
- runtime.display + barrel exports (display, display_protocol): a cached
`.display` client with info()/present(), the runtime.block shape.
- init spawns "display" in boot_services; build.zig wires the protocol onto the
runtime, builds the exe, packs it into the initial-ramdisk, installs it.
mmap fix the back buffer forced: systemMmap was capped at 256 pages (1 MiB) by a
fixed kernel-stack scratch array. Rewrote it to map page-by-page with rollback
(no array) and raised the cap to 8192 pages (32 MiB) — enough for a 4K back
buffer. A real limitation met.
Gate: `python3 test/qemu_test.py display-service` matches the service's own
serial heartbeats (display: online WxH / presented frame 0), printed only after
the full claim -> WC-map -> back-buffer -> present chain. Regression-checked
usermem, heap, init, and D1's display.
Kick off the display service track (docs/display.md, docs/display-plan.md): a
user-space compositor that owns the framebuffer. GOP and the PCI display device
are two views of one controller; GOP dies at ExitBootServices, so the portable
base is the boot-handoff linear framebuffer.
D1 makes that framebuffer reachable from user space over the existing device
claim/mmio_map path rather than a bespoke syscall:
- device-abi: a `display` DeviceClass, a DisplayInfo{w,h,pitch,format} on the
descriptor, and a flags field on resources with a write-combining bit.
- devices-broker: seedDisplay() publishes the loader's framebuffer as a
root-level `display` node (one WC-flagged memory resource); kmain seeds it
after discovery. displayDevice()/displayClaimed() track the claim.
- paging/mmio_map: mapUserDeviceInto gains a write_combining bool — a WC-flagged
resource maps through PAT entry 4 instead of strong-uncacheable (an
uncacheable framebuffer blit is glacial).
- console: falls silent while a display service holds the framebuffer, and is
forced back on by the panic/exception paths.
Gate: the `display` kernel test asserts the seeded node's shape and that the
claim + mmio_map leaf is genuinely write-combining (PAT bit set, PCD/PWT clear).
Regression-checked discovery/ioport/claim-release/supervision/device-list/
device-manager with the +1 device in the table.
init writes /mnt/usb/DANOS.LOG at shutdown and then enters S5 — but the write sat in the USB flash controller's write cache and was lost when power was cut, because nothing issued SCSI SYNCHRONIZE CACHE. Invisible in QEMU (its backing file commits immediately); real on hardware. The boot-time flush survived only because the machine kept running afterward and the cache drained on its own.
- block protocol gains a flush op; usb-storage serves it with SYNCHRONIZE CACHE (10); runtime.block gains Device.flush()
- the FAT server tracks whether blocks were written and, on a file close, commits the device cache (durable-on-close — the right default for removable media, and it makes init's existing shutdown close() persist the log before S5, no init/VFS change needed)
- verified the SYNCHRONIZE CACHE actually reaches the driver on each dirty close; usb-storage/fat-mount/mutations/rename/mtime/log-flush all green
Kernel discovery interpreted the whole DSDT/SSDTs (~0.5 MB -> ~9700 nodes) solely to extract the _S5 sleep type for a kernel-side soft-off — ~1-2s of work on the single-core, pre-scheduler critical path, with nothing to overlap it. The ring-3 acpi service already parses the same blobs and owns S5 end-to-end (reads pm1a_cnt from the FADT, writes SLP_TYP itself; orderly-shutdown tests it). So drop the kernel parse entirely.
- all *static*-table parsing stays (MADT/HPET/FADT/MCFG): CPUs, timers, PCIe, and the power register map are still detected from ACPI, not legacy/compat addresses (timer still calibrates via HPET/PM-timer/CPUID, PIT only as last resort)
- kernel keeps reboot (FADT reset register + legacy fallbacks — no AML); soft-off is userspace-only now
- removed PowerInformation.s5/s3, the kernel AML namespace, aml_stats, amlDeviceCount, and the aml import; power.shutdown/enable/sleepS3 gone
- tests: poweroff case retired (S5 covered by orderly-shutdown); discovery drops its S5/AML checks; acpi-parse self-verifies against a device-count floor since there's no kernel count to match
The console's one-time full-screen clear was millions of individual uncached word-writes to the GPU BAR: fine in QEMU, but on real hardware firmware MTRRs force that region uncacheable, so the clear crawls. Program PAT entry 4 to write-combining (setupPat, on the BSP and every AP) and map the framebuffer window with the PAT bit, so those writes batch into bursts.
The clear ran on the *loader's* uncached mapping because console.init happened before enablePaging. Move it to just after paging, so it uses the kernel's write-combining mapping instead. The cost: on-screen output is absent before paging (an early panic still lands in the serial/RAM log). Verified: 11 QEMU cases (incl. SMP AP-PAT, USB/FAT/ACPI) plus a screendump showing the framebuffer cleared to black with the status text rendered correctly.
The physmap (the kernel's permanent window onto all physical RAM) was built one 4 KiB page at a time: on a 64 GiB machine that is 16.7M mapPage calls and ~128 MiB of page tables (the bulk of the 'Kernel footprint' line). Map the 2 MiB-aligned interior with huge PD leaves instead — one entry per 2 MiB, no PT beneath — and only the unaligned head/tail with 4 KiB. 512x fewer entries and table frames; footprint at 8 GiB drops from ~16 MiB of tables to ~1 MiB total, and it scales.
translateIn/isExecutable now stop at a 2 MiB leaf, and descend() panics rather than walking through one (a 4 KiB map inside a huge page would corrupt RAM; the physmap and 4 KiB regions live in disjoint PML4 slots, so it can't happen — the guard just makes a bug loud). Verified: 20 QEMU cases (vmm/heap/wx/usermem/dma/ipc/smp/usb/fat/acpi/faults) green.
Serial is now a QEMU/dev aid, not a real-hardware necessity: a legacy-free board often has no live COM1, and the boot log is kept in RAM (klog) and flushed to disk. So the serial sink is compiled in only under -Dserial (default false).
- kernel.zig gates serialInit + the log sink on build_options.serial
- boot/efi.zig gates its EFI: progress breadcrumbs (con_out) via progress(); fatal-error messages stay always-on so a failed boot still explains itself
- run-x86-64 boots a serial-enabled image variant (factored addKernel/addBootImage helpers) so a dev boot always captures serial0, without baking serial into the flashable image
- test/qemu_test.py builds -Dserial=true (it asserts on serial markers)
- the loopback probe stays as a real-HW safety net for -Dserial images
On a legacy-free board COM1 is decoded but has no live UART: its LSR reads 0x00, so the transmit-holding-empty bit never sets and writeByte spun its full 100k-iteration guard on every logged byte (~15ms/byte x ~3.7k bytes to 'initialised' ~= 57s). QEMU always has a working UART, so this only bit real hardware; the boot log survived via log.ramSink.
- probe() loopback-tests the UART in init(); write() is a no-op when absent, so a dead port costs nothing per byte
- guard cut 100k -> 5k (backstop for a live-but-stalled UART only)
- serialPresent() + a boot line making the absent-UART case visible
Completes Phase 2: the FAT filesystem now stamps and reports a real modification
time, built on the Phase 2d(i) kernel wall-clock. This is the last stat field the
compiler's build cache needs to reason about (source vs cached output).
- on-disk.zig: fatToEpoch / epochToFatDateTime convert between the two 16-bit DOS
date/time fields and Unix epoch seconds (UTC — FAT has no timezone). Host-tested
round-trip + an absolute check (1577836800 == 2020-01-01).
- engine: a settable current_time_epoch that create/write stamp into the entry's
write (and creation) date/time; Node/Listing gained an mtime decoded from those
fields on read. Host test: a create stamps the mtime, read back through resolve
and listEntry.
- vfs protocol FileStatus + runtime.fs.Attributes gained an mtime field; the fat
server sets current_time_epoch from runtime.system.wallClock() per request and
returns mtime from stat. The flat ramfs reports 0 (it has no timestamps).
- fat-test reads the created file's mtime through stat and checks it is a real
current time, behind a new `fat-mtime` QEMU case.
Verified against the host: the guest stamped mtime 1783971676 while the host clock
was 1783971680 (boot+test lag) — the file's mtime is real current time. zig build,
zig build test (the epoch<->DOS conversions + the engine mtime test),
zig build check-fat-image, and a sequential QEMU sweep — fat-mount, fat-mutations,
fat-rename, fat-mtime, vfs, vfs-client-death, log-flush, orderly-shutdown,
initial-ramdisk, smoke, wall-clock, usb-storage — all green. mode/inode remain.
Adds real (calendar) time, the foundation for filesystem timestamps. Monotonic
time (`clock`) says how long since boot; this says what time it actually is.
- cpu.zig (x86_64): readRtcUnixSeconds() reads the CMOS real-time clock (ports
0x70/0x71) — waits out an update-in-progress, reads twice until stable, handles
BCD-vs-binary and 12-vs-24-hour per status register B — and converts to Unix
epoch seconds (UTC).
- kernel/wall-clock.zig: reads the RTC once at boot and anchors it to the monotonic
clock, so a query is a cheap arithmetic offset — no per-call CMOS poll, no lock,
no SMP hazard on the shared ports. kmain calls init() once the monotonic clock is
final and logs the epoch.
- wall_clock() syscall (33) -> Unix epoch seconds, wrapped by runtime.system
.wallClock(). Wall-clock *seconds* are mechanism the kernel owns like the
monotonic clock; calendars/timezones are user-space policy (the stale comment on
systemClock that called wall-clock a "user-space service" is updated in spirit by
the new handler's doc).
- A `wall-clock` kernel test asserts the boot RTC read is a plausible current epoch.
Verified against the host: the guest read epoch 1783971244 while `date -u +%s` gave
1783971245 (one second of boot lag) — the CMOS read + epoch conversion are correct to
the second. zig build, zig build test, and smoke/clock/init are green.
Completes the Phase 2 FAT mutation set (truncate, mkdir, unlink, rename).
- engine: rename(dir, old_name, new_name) rewrites an existing entry's 8.3 name in
place within the same directory. Refuses a missing source, a non-8.3 target, or a
name that already exists; drops any long-name entries on the old file (it takes
its new 8.3 name), LFN-aware like removeFile. Cross-directory and long-name-
preserving rename are noted limitations. Host-tested (rename keeps contents;
collision, non-8.3, and missing-source are refused).
- vfs protocol: a `rename` operation whose payload is old-path, a 0x00 separator,
then new-path.
- VFS router: a forwardRename helper + a `.rename` case that requires both paths
under the same mount (cross-filesystem rename is refused) and forwards the
mount-relative old+new.
- fat server: a `.rename` handler that requires the same parent directory and calls
engine.rename.
- runtime.fs: rename(old_path, new_path).
- fat-test now renames the file it created (before removing it) and asserts the old
name is gone, behind a new `fat-rename` QEMU case.
Verified: zig build, zig build test (the engine rename unit test), zig build
check-fat-image, and a sequential QEMU sweep — fat-mount, fat-mutations, fat-rename,
vfs, vfs-client-death, log-flush, orderly-shutdown, initial-ramdisk, smoke — green.
The engine gained mkdir/unlink in Phase 2a; this exposes them as first-class
filesystem operations so programs can use them.
- vfs protocol: two new path-based operations, mkdir and unlink (appended, so
existing opcodes/offsets are unchanged).
- VFS router: a forwardPath helper relays a path-based op under a mount to its
backend; the mkdir/unlink cases forward to the mounted filesystem (the flat
ramfs refuses them — it has no directories).
- fat server: mkdir -> engine.createDirectory, unlink -> engine.removeFile, each
resolving the parent via a shared splitParent helper (also used by open-create).
- runtime.fs: makeDirectory(path) and remove(path).
- fat-test now exercises the whole path — mkdir /mnt/usb/TESTDIR, create + write +
read a file inside it, then remove it — behind a new `fat-mutations` QEMU case.
Verified: zig build, zig build test, zig build check-fat-image, and a sequential
QEMU sweep — fat-mount, fat-mutations (mkdir/write/read/unlink through the mount),
vfs, vfs-client-death, log-flush, orderly-shutdown, initial-ramdisk, smoke — green.
The FAT engine could create, read, write, and grow files, but never free clusters
or make directories — so overwriting a shorter file left a stale tail (a real bug:
corrupt boot-log re-flushes, and later corrupt compiler cache/.o files). This adds
the mutation half of the engine, with the corruption fix wired all the way through.
Engine (system/services/fat/engine.zig), all host-tested:
- freeChain: return a cluster chain to the pool (bounded against a corrupt cycle) —
the shared primitive under truncate and remove.
- truncate: free the chain and zero the entry's size/first-cluster (O_TRUNC).
- createDirectory (mkdir): allocate + initialise a cluster with "." and ".." and add
the directory entry to the parent.
- removeFile (unlink): free the chain and mark the 8.3 entry plus any preceding
long-name entries deleted, so a reused slot can't inherit an orphaned long name.
Refuses directories.
- createFile now shares a common addEntry helper with createDirectory.
O_TRUNC wired end to end: a truncate open-flag (vfs protocol) that the router already
forwards; runtime.fs.OpenOptions.truncate; and fat's handleOpen calls engine.truncate
on an existing file. The boot-log flush (log-flush + init) now opens with truncate, so
a shorter log on a later boot of the same stick leaves no stale tail — closing the
caveat from the boot-log work.
mkdir/unlink are engine-complete and host-tested but not yet exposed as VFS
operations / runtime.fs methods (they are new path-based ops needing router cases);
rename and the richer stat (mtime/mode, blocked on wall-clock) remain. See
docs/zig-self-hosting.md (Phase 2).
Verified: zig build, zig build test (8 engine host tests, incl. truncate, the
overwrite-no-stale-tail regression, remove, and mkdir), zig build check-fat-image, and
a sequential QEMU sweep — fat-mount, log-flush (DANOS.LOG read back at 11804 bytes,
clean, with the truncate-based final flush), vfs, vfs-client-death, usb-storage,
orderly-shutdown, initial-ramdisk, smoke — all green.
The first step of the Zig self-hosting roadmap (docs/zig-self-hosting.md): give
danos programs a danos-native file API and remove the premature POSIX compatibility
shim. This also resolves the earlier misplacement of a full-write helper into the
compat layer — that behaviour now lives natively in runtime.fs.File.writeAll.
- library/runtime/fs.zig: the danos-native file client over the VFS (open/read/
write/writeAll/seekTo/attributes/close, directory listing, mount). Handles are
*values* — a File/Directory owns its VFS node id and byte offset — so there is no
per-process fd table or descriptor limit, unlike the POSIX fd model the shim
emulated. This is where the operations that later become std.os.danos are staged.
- Retire library/posix/ (unistd, stdio): only five call sites used it, all file
operations, all migrated to runtime.fs — fat (mount), the vfs-test and fat-test
clients, and init/log-flush (the boot-log flush). stdio was already dead.
- build.zig: drop the posix module, its addUserBinary parameter, the per-binary
import, and the ~26 call-site arguments.
- Docs: the VFS protocol's client is now runtime.fs; the docs index and
coding-standards note posix is retired and the foreign-ABI naming exception now
applies to the future std.os.danos seam; the process-lifecycle note points the
future musl layer at that same seam rather than the deleted directory.
Deferred by design (see the roadmap): the C-ABI runtime.os errno seam is built at
fork time (its shape must match std/os/danos.zig); truncate/mkdir/rename are
Phase 2; stdio-byte fds and cwd are later slices.
Verified: zig build, zig build test, zig build check-fat-image, and a sequential
QEMU sweep — vfs, vfs-client-death (the park/hold-handle path), fat-mount, log-flush,
orderly-shutdown, initial-ramdisk (log-flush silent in the bare sweep), smoke, init,
usb-storage, device-manager — all green.
On a headless or real board nothing captures serial, so the boot log — the whole
diagnostic stream — is lost at power-off. This retains it in the kernel and copies
it to the boot USB volume as /mnt/usb/DANOS.LOG, the on-disk equivalent of QEMU's
`-serial file:`. Pull the stick, read DANOS.LOG on another machine.
How it fits together:
- Kernel RAM sink (log.zig): a fixed 256 KiB in-image buffer registered as a log
sink in kmain, right after serial. Because userspace debug_write funnels through
log.write, it captures the entire stream — kernel lines and every service's
output — from the first line. Fills linearly and stops when full (earliest boot
output, the most valuable, is kept); no allocation, so it is panic-safe.
- klog_read syscall (32): copies that buffer out to a user buffer, the mirror of
debug_write — same overflow-safe user-half bounds check, kernel -> user copy,
under the kernel lock so the snapshot can't grow mid-copy. Wrapped by
runtime.system.klogRead.
- log-flush (new one-shot, in the initial-ramdisk): waits for the fat server to
mount /mnt/usb, then copies the whole log to /mnt/usb/DANOS.LOG. init spawns it
once the boot services are up (fire-and-forget; it polls the mount itself). If
no volume is mounted — no stick, or the no-VFS ramdisk sweep — it exits silently.
- init shutdown flush: init repeats the copy inline at the top of shutDown(),
BEFORE it tears down the storage services (the fat server is stopped first), so
a clean poweroff captures the fullest log while /mnt/usb is still writable.
- unistd.writeAll: loops write() past the 224-byte VFS payload cap; both flush
paths use it.
The filename is 8.3 (DANOS.LOG) at the mount root — the FAT short-name rule, and
there is no mkdir on the FAT path yet. Extend-only writes mean the two same-session
flushes never leave stale bytes (the shutdown log is a superset of the boot log);
a shorter log on a later boot of the same stick can leave a stale tail — a noted,
cosmetic limitation, not worth pulling O_TRUNC into the FAT write path for now.
Verified end to end under QEMU: a new `log-flush` case (reusing the orderly-shutdown
build) asserts both markers then S5, and DANOS.LOG is read back out of the image
afterwards (11804 bytes, containing the kernel init line, the FAT mount line, and
the boot flush marker). Regression stays green: zig build, zig build test, zig
build check-fat-image, and a sequential QEMU sweep — smoke, init, initial-ramdisk
(log-flush silent in the bare sweep), orderly-shutdown, fat-mount, usb-storage,
usb-hid, vfs, input, device-manager, process, signals, dma, fault-pf.
Add a FAT12/16/32 filesystem the VFS mounts at /mnt/usb, reading and writing a
USB stick through the block device. Verified end to end under QEMU: the fat
server mounts the volume, the VFS routes /mnt/usb to it, and a client lists the
root and reads a file (the ELF magic of /mnt/usb/system/kernel).
- engine.zig: the FAT engine over a BlockDevice interface — mount (a bare FAT or,
as QEMU's VVFAT and most real sticks present it, an MBR-partitioned disk), FAT
chain walk (12/16/32), cluster allocation, directory traversal with long-name
read, and file read / write / create. Host-tested against a RAM-backed FAT16
image (create, cluster-spanning write, mid-file overwrite, read-back, list).
- on-disk.zig: the align(1) boot-sector / directory / long-name / FSInfo structs
and the cluster-count FAT-type detection.
- fat.zig: the server — wraps the .block device (a DMA bounce buffer) in a
BlockDevice, mounts the FAT, serves the vfs-protocol as a backend, and mounts
itself into the VFS at /mnt/usb. Spawned by init as a boot service.
- runtime.block: the block-device client (geometry / read / write by physical
address, so whole sectors never cross IPC).
- Raise the kernel service-name registry (maximum_services) 8 -> 16: it is
indexed directly by ServiceId, and fat = 8 was being rejected, so the fat
server exited before registering.
- VFS: an absolute path with no matching mount is now not-found rather than
silently created in the flat ramfs — so /mnt/usb fails cleanly until mounted.
Tests: fat-mount (the full stack: block -> FAT -> VFS mount -> list + file read)
passes; host units cover the engine and on-disk structs; the vfs, shutdown, and
USB regression suite stays green (10/10).
Daniel Samson
Daniel Samson