Process lifecycle and state¶

Scope¶

ChrisOS process management in the reviewed revision is deliberately bounded and explicit. There are at most 32 process slots. Slot zero is the kernel. User processes own a page-table root, a bounded table of owned user pages, fixed-region virtual-memory metadata, a small state value and a short name. User process switching occurs only on the BSP.

This is not a conventional mature scheduler with arbitrary threads per process, per-CPU run queues and signal semantics. It is nevertheless a real process lifecycle: creation allocates an address space and initial stack, mapping establishes owned physical frames, page faults can commit memory lazily, blocking changes runnable state, destruction tears down frames/address space and closes resources.

Process table¶

g_proc[PROC_MAX] is a static array with PROC_MAX = 32.

The internal Proc stores: - used: slot allocation; - alive: whether execution is still valid; - state: FREE, READY or one of the blocking reasons; - cr3: process page-table root; - vm_bytes: length of the VM demand region; - heap_brk: current heap break; - npages: count of owned pages; - fb_pages: framebuffer-region page count; - 24-byte process name; - up to 288 ProcPage ownership records.

Each ProcPage records virtual and physical address for a frame owned by the process.

This ownership table is central to teardown: mappings alone do not tell the kernel which physical frames it should return to PMM.

Fixed states¶

The public header defines: - PROC_ST_FREE; - PROC_ST_READY; - PROC_ST_BLOCK_SOCK; - PROC_ST_BLOCK_JOIN; - PROC_ST_BLOCK_IRQ.

A process is considered runnable only when its slot is valid, alive and state is READY.

The state is intentionally small. There is no separate zombie state, stop/continue state, priority or per-thread state in this structure.

Initialization¶

proc_init resets every slot: not used, not alive, FREE, no CR3, no VM allocation, heap break reset to PROC_HEAP_VIRT, no pages and empty name.

Then it installs slot zero as the kernel: - used = 1; - alive = 1; - state = READY; - CR3 = mm_kernel_cr3(); - name = “kernel”.

g_current becomes PROC_KERNEL and the slice flag and last-fault record are reset.

The kernel process therefore participates in the same current-PID namespace even though it does not use the user-page ownership path.

Creating a process¶

proc_create(name) scans slots 1–31 for an unused entry.

For a candidate it calls mm_clone_kernel_space(). A zero return aborts creation.

The slot is then initialized as alive/READY, with its new CR3, empty owned-page table, heap break at PROC_HEAP_VIRT and copied name.

Creation immediately calls:

proc_commit(i, PROC_STACK_VIRT)

If initial stack-page commitment fails, proc_destroy(i) unwinds the process and creation returns -1.

A successfully returned PID therefore already has an address space and at least one owned user stack page.

Kernel mappings in a process¶

The name mm_clone_kernel_space indicates that user address spaces retain the kernel mappings needed for controlled entry/exit while obtaining an independent lower user space. Exact page-table topology belongs to the virtual-memory chapter.

The process object stores the returned CR3 rather than copying kernel virtual mappings itself.

This design means the privilege boundary is enforced by page flags, not by making kernel code completely absent from every CR3. Supervisor mappings may exist while CPL 3 is forbidden from accessing them.

User virtual layout¶

Current constants are:

PROC_PAGES = 288 bounds the number of tracked owned pages. The VM byte limit is clamped to that many pages. Heap growth is separately limited to 256 KiB.

These are concrete capacity choices, not general x86 limits.

Mapping without and with ownership¶

proc_map_user validates the PID and lower-half address, then invokes mm_map_cr3 with MM_USER added to supplied flags.

proc_map_owned additionally: - page-aligns the virtual address; - rejects duplicates; - rejects ownership-table overflow; - maps the frame; - records virtual/physical pair.

On failure to record/map, the caller remains responsible for freeing a frame that was not transferred into process ownership.

The distinction between “map” and “map + own” prevents ambiguous teardown responsibility.

Committing a page¶

proc_commit is the standard owned-page allocation path.

It aligns the address and treats an already-owned page as success. It checks the 288-entry limit, allocates a physical frame with pmm_alloc and zero-fills the entire 4 KiB page through the physical-to-virtual mapping.

It maps the frame as Present + Write + User. If mapping fails, the frame is returned to PMM.

Then it appends the ownership record. If the process is currently selected, the function flushes TLB.

Zero-fill is both a correctness and security property: a newly exposed user page must not reveal data left by a previous frame owner.

Demand regions¶

proc_set_vm(pid, bytes) clamps VM size to the process-page capacity, stores the length and eagerly commits the first VM page. The rest can arrive on demand.

proc_fault_demand page-aligns CR2 and recognizes: - VM range from PROC_VM_VIRT to PROC_VM_VIRT + vm_bytes; - four stack pages starting at PROC_STACK_VIRT; - heap pages below current heap_brk; - framebuffer pages within configured count.

If CR2 lies in one of these regions, it calls proc_commit. Otherwise it reports that the fault is not recoverable by demand allocation.

This turns region metadata into a page-fault policy.

Heap growth¶

proc_sbrk(pid, inc) returns the old break and advances the process break if the result stays within 256 KiB above PROC_HEAP_VIRT.

It does not eagerly allocate every page crossed. Later access faults and proc_fault_demand commits pages below the new break.

The current function also does not implement negative increments or heap shrink. Its unsigned interface is grow-only.

Framebuffer region¶

proc_fb_ptr(pid, pages) clamps requested pages to at most 16, stores the count and eagerly commits each framebuffer-region page.

It returns the fixed virtual base PROC_FB_VIRT.

This is a process-owned software buffer mapping, not direct ownership of a physical GPU framebuffer. The graphics/security implications should be distinguished from MMIO mapping.

Switching¶

proc_switch(pid) first enforces:

if (smp_current_cpu() != 0u)
    panic("process switch off BSP");

For a valid used PID it reads the process CR3, updates g_current and calls mm_switch.

The invariant is explicit: user process scheduler state is global and is not safe to mutate from APs.

This rule prevents simultaneous CPUs from treating one global g_current as their independent current process.

Timer slice state¶

The PIT handler eventually calls proc_on_tick, which sets g_slice = 1.

proc_slice_due reads this flag; proc_slice_ack clears it.

This is scheduling metadata, not the complete process switch. The lifecycle code does not contain a fully preemptive multi-process context-switch state machine.

Blocking¶

proc_block(pid, why) replaces process state with a blocking reason. proc_unblock(pid) restores READY. proc_unblock_why(why) scans every user process and makes all processes blocked for that reason READY.

The API is intentionally coarse. It does not maintain per-object waiter queues in the process table itself.

Subsystems such as sockets or join semantics can use these states as a global scheduling signal.

Fault records and liveness¶

proc_record_fault stores one global ProcFault containing valid flag, pid, tid, CR2 and RIP. If the PID is valid it sets alive = 0 and logs identifying information.

proc_last_fault exposes the record.

The record is “last fault” diagnostic state, not a per-process history. A later fault overwrites it.

User-fault handling then destroys the process and arranges return to kernel code.

Destruction¶

proc_destroy(pid) rejects kernel PID and invalid slots.

If the process being destroyed is current, it first switches to PROC_KERNEL. This is essential: freeing the active user page tables while CR3 still points at them would destroy the address-space context executing teardown.

It then: 1. releases every owned user frame; 2. frees the user page-table structure if it is distinct from kernel CR3; 3. closes syscall-owned file descriptors for the PID; 4. closes sockets owned by the process; 5. marks the slot FREE/not alive and clears CR3.

This is ordered teardown: leave the address space before dismantling it, release mappings/frames, release subsystem resources, then make the slot reusable.

Releasing user pages¶

proc_release_user iterates ownership records. For each physical frame: - unmaps the virtual address from process CR3 if a CR3 remains; - returns physical frame to PMM; - clears record fields.

After the loop it resets page count, VM bytes, framebuffer page count and heap break.

The ownership table ensures each recorded frame is freed once. Duplicate virtual addresses are rejected during owned mapping, reducing double-ownership risk.

File and socket cleanup¶

syscall_close_owner(pid) scans the small syscall file table and invalidates descriptors belonging to the process.

sock_close_proc(pid) performs analogous network-resource cleanup in the socket layer.

This demonstrates that process destruction is a cross-subsystem operation. Adding future resource classes requires either extending teardown or introducing a more general owned-handle mechanism.

Current-process destruction¶

The initial switch to kernel PID before release also changes CR3 through the normal proc_switch path.

Because only the BSP is allowed to switch processes, destruction of the current user process is expected on BSP. A future concurrent architecture must guarantee that no other CPU can still execute in the address space before pages/page tables are freed.

This requirement connects process teardown directly to TLB shootdown and task migration.

Capacity and failure¶

Creation fails if: - no free process slot exists; - kernel-space cloning fails; - initial stack commit fails.

Commit fails if: - PID invalid; - ownership table full; - PMM allocation fails; - page mapping fails.

Framebuffer setup can partially commit pages before a later failure; callers need to treat the process-owned page table as the source of teardown truth.

No overcommit policy or swap exists in the reviewed path.

Security properties¶

User frames are zeroed before exposure. Mappings receive MM_USER only through the user mapping path. Virtual addresses at or above the lower canonical limit are rejected in proc_map_user.

Destroying a process closes its user-visible file/socket ownership, limiting handle leakage into a reused slot.

The fixed PID slot is not generation-tagged. If external code stores a PID after destruction and later the slot is reused, stale PID references could refer to a different process unless callers enforce lifetime. A mature handle model commonly pairs index with generation or stronger references.

Performance characteristics¶

Process lookup and creation scan a fixed 32-entry array: bounded O(32). Owned-page lookup is linear in at most 288 records. Destruction similarly walks the owned list.

For current small limits this favors simplicity over sophisticated VM trees. Scaling to thousands of mappings/processes would require different data structures.

Demand paging reduces eager allocation but adds first-touch faults.

Validation¶

Lifecycle tests should verify: - all 31 user slots and exhaustion; - clone failure/stack allocation unwind; - duplicate owned mapping rejection; - zero-filled new pages; - VM/stack/heap/framebuffer demand boundaries; - heap maximum; - switch-on-AP panic invariant; - block/unblock transitions; - destruction of current process switches to kernel first; - every owned frame returns exactly once; - file/socket ownership closes; - PID slot can be reused only after complete cleanup.

A leak test comparing PMM free state before/after repeated create/destroy cycles is especially valuable.

Current limitations¶

One process-level execution context is represented per PID; no user threads are modeled here. Current process and scheduling state are global/BSP-only. Limits are static. Heap only grows. Fault history is one global record. Resource teardown is explicitly wired to known subsystems.

These limits should remain visible instead of being obscured by generic language such as “full process management”.

Source map¶

kernel/metal/proc.c/proc.h define process state and lifecycle. mm.c/mm.h provide CR3/mapping operations. pmm.c/pmm.h own physical frames. syscall.c and kernel/net/sock.c participate in teardown. The Source Atlas publishes each source in full at revision da3df29cb397932c43d32373871fb9380e688ade.