Memory Management

minix.rs splits memory management the way MINIX 3 does — mechanism in the kernel, policy in a user-space server — but moves physical-frame ownership into the kernel. The kernel owns the frame allocator and performs every page table write; the user-space VM server decides what should be mapped where and drives the kernel through a single privileged kernel call (SYS_VMCTL).

This chapter describes the subsystem as it stands at the end of Phase 3: a physical frame allocator, per-process address spaces with hardware-tagged TLBs, a page-fault path that delegates to VM, and the VM server’s region tracking for brk, mmap, and munmap.

Physical frame allocator

kernel/src/mm/ is the kernel-side physical allocator. At boot it walks Limine’s MEMMAP_USABLE entries and seeds a per-region bump pointer for each (MAX_REGIONS = 16 regions — QEMU virt and Apple-Silicon QEMU both fit comfortably). Freed frames go onto an intrusive free-list threaded through the frames themselves, reached via Limine’s higher-half direct map (HHDM).

Two invariants keep callers honest:

• alloc_frame zeroes every frame before handing it out, so a caller never observes residual state.
• free_frame pushes the frame back through the HHDM.

Frames inside the kernel image, the embedded boot image, and the static EL0 test-stub pages live in Limine’s EXECUTABLE_AND_MODULES region and are never visible to the allocator, so no explicit reservation logic is needed.

Per-process address spaces

Each user process has its own translation table tree, rooted at a physical address recorded in Proc::ttbr0_pa, and an 8-bit address-space identifier in Proc::asid. kernel/src/arch/aarch64/addrspace.rs is the page-table API:

• AddrSpace::new() allocates an L0 root frame.
• map_page(va, pa, prot) walks the tree, allocating L1/L2/L3 tables on demand from the frame allocator, and writes the leaf PTE through the HHDM.
• walk_pt(va) resolves a VA to its PTE (or None).
• destroy() recursively frees the intermediate tables and the L0 root (leaf frames are caller-owned and freed elsewhere).

The free functions map_page_in(ttbr0_pa, …) and unmap_page_in(ttbr0_pa, …) do the same work keyed by a root PA, so the kernel can mutate a process’s tree without holding an AddrSpace value — this is how SYS_VMCTL edits a target process’s address space.

Context switch and ASIDs

On every switch into a user process, proc::sched::schedule_next:

1. parks the next process’s register frame for the trap return,
2. calls switch_ttbr0_with_asid(ttbr0_pa, asid) — which writes TTBR0_EL1 = ttbr0_pa | ((asid as u64) << 48) and issues an ASID-tagged tlbi — before
3. flushing any pending IPC message into the user buffer.

The order matters: the message flush writes through the active TTBR0, so the incoming process’s address space must be live first. ASIDs let the hardware keep TLB entries for multiple address spaces without a full flush on each switch; the allocator (asid.rs) hands them out from FIRST_ASID = 1 (0 means “uninitialized”) and panics on 8-bit wrap — real rollover is deferred until process churn in Phase 4 makes it reachable.

The page-fault path

When an EL0 process touches an unmapped page, the CPU traps to EL1 and do_page_fault(esr, elr, far) runs. It classifies the abort (instruction vs data, fault status code, write-vs-read), records the coordinates in the faulting process’s PageFaultState, blocks the process on the RTS_PAGEFAULT run-time state, and sends the VM server a VM_PAGEFAULT message carrying the faulting endpoint and the fault address. Permission faults (a write to a read-only page) are not resolvable by mapping and halt loudly.

The kernel originates that send with mini_pf_send, which models the faulting process as a blocked sender on VM’s caller queue — the lingering RTS_PAGEFAULT keeps it blocked even after the RTS_SENDING half clears, until VM explicitly clears the fault. Because VM resolves the fault while running under its own TTBR0 and the message is delivered after the switch, no cross-address-space copy machinery is needed at this stage.

SYS_VMCTL: kernel mechanism, VM policy

SYS_VMCTL (kernel/src/system/do_vmctl.rs) is the one privileged call through which VM drives the kernel’s paging mechanism. It is gated solely by Priv::k_call_mask granting SYS_VMCTL; VM is its sole intended holder and is trusted to target only processes it legitimately manages (the MINIX 3 trust model). Each subcall names a target process by endpoint (SELF allowed):

Every PTE change is followed by an ASID-tagged TLB invalidation. The kernel allocates frames (unlike MINIX 3, where VM owns physical memory) and VM supplies only the virtual address and protection.

VM server region tracking

The VM server (servers/vm/) is the first real user-space process. It runs a RECEIVE(ANY) loop and dispatches on message type. It owns no heap allocator — the kernel owns frames — so it tracks memory with a static per-process region table (servers/vm/src/region.rs): [ClientRegions; 16], keyed by process number, each holding up to MAX_REGIONS = 4 regions. A region is a half-open virtual range [start, end) tagged with a Kind:

• Heap — grown by brk, based at the fixed HEAP_BASE (0x0100_0000).
• Mmap — anonymous mappings, bump-allocated from MMAP_BASE (0x0200_0000).
• Unused — a free slot.

A page fault is satisfied only if its address lies inside one of the faulting process’s regions: VM consults the table, and on a hit issues SYS_VMCTL(VMCTL_PT_MAP) then SYS_VMCTL(VMCTL_CLEAR_PAGEFAULT). A fault outside every region is a SIGSEGV — VM leaves the process blocked on RTS_PAGEFAULT (the only “kill” available until PM and signals arrive in Phase 4). VM runs at EL0 with no console, so this path is silent; the symptom is a process that stops making progress.

brk

VM_BRK(new_break) sets the caller’s program break. VM page-aligns the request and grows (or first creates) the caller’s Heap region to [HEAP_BASE, new_break), replying with the resulting break. No frames are mapped eagerly — pages fault in lazily on first touch and are then satisfied by the region check above.

mmap / munmap

VM_MMAP(len) is an anonymous mapping in the style of mmap(NULL, len, …): VM page-aligns the length, bump-allocates a base address from the caller’s mmap arena, records an Mmap region, and replies with the chosen base. As with the heap, frames are not mapped until first touch.

VM_MUNMAP(addr, len) drops the Mmap region based at addr and unmaps each backing page with SYS_VMCTL(VMCTL_PT_UNMAP). Pages that never faulted in were never mapped, so the kernel returns a harmless EINVAL for them, which VM ignores. The match is keyed on the region’s base address and the unmap sweep is capped at the region’s own end, so an over-stated length can never reach into a neighboring region or the heap.

The arena is bump-only for now — munmap does not return addresses to it. Real address reuse and a PM-supplied per-process memory layout arrive in Phase 4.