In the previous posts we created a frame allocator and a page table module. Now we are ready to create a kernel heap and a memory allocator. Thus, we will unlock Box, Vec, BTreeMap, and the rest of the alloc crate.

As always, you can find the complete source code on GitHub. Please file issues for any problems, questions, or improvement suggestions. There is also a comment section at the end of this page.

Introduction

The heap is the memory area for long-lived allocations. The programmer can access it by using types like Box or Vec. Behind the scenes, the compiler manages that memory by inserting calls to some memory allocator. By default, Rust links to the jemalloc allocator (for binaries) or the system allocator (for libraries). However, both rely on system calls such as sbrk and are thus unusable in our kernel. So we need to create and link our own allocator.

A good allocator is fast and reliable. It also effectively utilizes the available memory and keeps fragmentation low. Furthermore, it works well for concurrent applications and scales to any number of processors. It even optimizes the memory layout with respect to the CPU caches to improve cache locality and avoid false sharing.

These requirements make good allocators pretty complex. For example, jemalloc has over 30.000 lines of code. This complexity is out of scope for our kernel, so we will create a much simpler allocator. Nevertheless, it should suffice for the foreseeable future, since we’ll allocate only when it’s absolutely necessary.

The Allocator Interface

The allocator interface in Rust is defined through the Alloc trait, which looks like this:

pub unsafe trait Alloc {
    unsafe fn alloc(&mut self, layout: Layout) -> Result<*mut u8, AllocErr>;
    unsafe fn dealloc(&mut self, ptr: *mut u8, layout: Layout);
    […] // about 13 methods with default implementations
}

The alloc method should allocate a memory block with the size and alignment given through Layout parameter. The deallocate method should free such memory blocks again. Both methods are unsafe, as is the trait itself. This has different reasons:

• Implementing the Alloc trait is unsafe, because the implementation must satisfy a set of contracts. Among other things, pointers returned by alloc must point to valid memory and adhere to the Layout requirements.
• Calling alloc is unsafe because the caller must ensure that the passed layout does not have size zero. I think this is because of compatibility reasons with existing C-allocators, where zero-sized allocations are undefined behavior.
• Calling dealloc is unsafe because the caller must guarantee that the passed parameters adhere to the contract. For example, ptr must denote a valid memory block allocated via this allocator.

To set the system allocator, the global_allocator attribute can be added to a static that implements Alloc for a shared reference of itself. For example:

#[global_allocator]
static MY_ALLOCATOR: MyAllocator = MyAllocator {...};
impl<'a> Alloc for &'a MyAllocator {
    unsafe fn alloc(&mut self, layout: Layout) -> Result<*mut u8, AllocErr> {...}
    unsafe fn dealloc(&mut self, ptr: *mut u8, layout: Layout) {...}
}

Note that Alloc needs to be implemented for &MyAllocator, not for MyAllocator. The reason is that the alloc and dealloc methods require mutable self references, but there’s no way to get such a reference safely from a static. By requiring implementations for &MyAllocator, the global allocator interface avoids this problem and pushes the burden of synchronization onto the user.

Including the alloc crate

The Alloc trait is part of the alloc crate, which like core is a subset of Rust’s standard library. Apart from the trait, the crate also contains the standard types that require allocations such as Box, Vec and Arc. We can include it through a simple extern crate:

// in src/lib.rs
#![feature(alloc)] // the alloc crate is still unstable
[...]
#[macro_use]
extern crate alloc;

We don’t need to add anything to our Cargo.toml, since the alloc crate is part of the standard library and shipped with the Rust compiler. The alloc crate provides the format! and vec! macros, so we use #[macro_use] to import them.

When we try to compile our crate now, the following error occurs:

error[E0463]: can't find crate for `alloc`
  --> src/lib.rs:10:1
   |
16 | extern crate alloc;
   | ^^^^^^^^^^^^^^^^^^^ can't find crate

The problem is that xargo only cross compiles libcore by default. To also cross compile the alloc crate, we need to create a file named Xargo.toml in our project root (right next to the Cargo.toml) with the following content:

[target.x86_64-blog_os.dependencies]
alloc = {}

This instructs xargo that we also need alloc. It still doesn’t compile, since we need to define a global allocator in order to use the alloc crate:

error: no #[default_lib_allocator] found but one is required; is libstd not linked?

A Bump Allocator

For our first allocator, we start simple. We create a memory::heap_allocator module containing a so-called bump allocator:

// in src/memory/mod.rs
mod heap_allocator;
// in src/memory/heap_allocator.rs
use alloc::heap::{Alloc, AllocErr, Layout};
/// A simple allocator that allocates memory linearly and ignores freed memory.
#[derive(Debug)]
pub struct BumpAllocator {
    heap_start: usize,
    heap_end: usize,
    next: usize,
}
impl BumpAllocator {
    pub const fn new(heap_start: usize, heap_end: usize) -> Self {
        Self { heap_start, heap_end, next: heap_start }
    }
}
unsafe impl Alloc for BumpAllocator {
    unsafe fn alloc(&mut self, layout: Layout) -> Result<*mut u8, AllocErr> {
        let alloc_start = align_up(self.next, layout.align());
        let alloc_end = alloc_start.saturating_add(layout.size());
        if alloc_end <= self.heap_end {
            self.next = alloc_end;
            Ok(alloc_start as *mut u8)
        } else {
            Err(AllocErr::Exhausted{ request: layout })
        }
    }
    unsafe fn dealloc(&mut self, ptr: *mut u8, layout: Layout) {
        // do nothing, leak memory
    }
}

We also need to add #![feature(allocator_api)] to our lib.rs, since the allocator API is still unstable.

The heap_start and heap_end fields contain the start and end address of our kernel heap. The next field contains the next free address and is increased after every allocation. To allocate a memory block we align the next address using the align_up function (described below). Then we add up the desired size and make sure that we don’t exceed the end of the heap. We use a saturating add so that the alloc_end cannot overflow, which could lead to an invalid allocation. If everything goes well, we update the next address and return a pointer to the start address of the allocation. Else, we return None.

Alignment

In order to simplify alignment, we add align_down and align_up functions:

/// Align downwards. Returns the greatest x with alignment `align`
/// so that x <= addr. The alignment must be a power of 2.
pub fn align_down(addr: usize, align: usize) -> usize {
    if align.is_power_of_two() {
        addr & !(align - 1)
    } else if align == 0 {
        addr
    } else {
        panic!("`align` must be a power of 2");
    }
}
/// Align upwards. Returns the smallest x with alignment `align`
/// so that x >= addr. The alignment must be a power of 2.
pub fn align_up(addr: usize, align: usize) -> usize {
    align_down(addr + align - 1, align)
}

Let’s start with align_down: If the alignment is a valid power of two (i.e. in {1,2,4,8,…}), we use some bitwise operations to return the aligned address. It works because every power of two has exactly one bit set in its binary representation. For example, the numbers {1,2,4,8,…} are {1,10,100,1000,…} in binary. By subtracting 1 we get {0,01,011,0111,…}. These binary numbers have a 1 at exactly the positions that need to be zeroed in addr. For example, the last 3 bits need to be zeroed for a alignment of 8.

To align addr, we create a bitmask from align-1. We want a 0 at the position of each 1, so we invert it using !. After that, the binary numbers look like this: {…11111,…11110,…11100,…11000,…}. Finally, we zero the correct bits using a binary AND.

Aligning upwards is simple now. We just increase addr by align-1 and call align_down. We add align-1 instead of align because we would otherwise waste align bytes for already aligned addresses.

Reusing Freed Memory

The heap memory is limited, so we should reuse freed memory for new allocations. This sounds simple, but is not so easy in practice since allocations can live arbitrarily long (and can be freed in an arbitrary order). This means that we need some kind of data structure to keep track of which memory areas are free and which are in use. This data structure should be very optimized since it causes overheads in both space (i.e. it needs backing memory) and time (i.e. accessing and organizing it needs CPU cycles).

Our bump allocator only keeps track of the next free memory address, which doesn’t suffice to keep track of freed memory areas. So our only choice is to ignore deallocations and leak the corresponding memory. Thus our allocator quickly runs out of memory in a real system, but it suffices for simple testing. Later in this post, we will introduce a better allocator that does not leak freed memory.

Using it as System Allocator

Above we saw that we can use a static allocator as system allocator through the global_allocator attribute:

#[global_allocator]
static ALLOCATOR: MyAllocator = MyAllocator {...};

This requires an implementation of Alloc for &MyAllocator, i.e. a shared reference. If we try to add such an implementation for our bump allocator (unsafe impl<'a> Alloc for &'a BumpAllocator), we have a problem: Our alloc method requires updating the next field, which is not possible for a shared reference.

One solution could be to put the bump allocator behind a Mutex and wrap it into a new type, for which we can implement Alloc for a shared reference:

struct LockedBumpAllocator(Mutex<BumpAllocator>);
impl<'a> Alloc for &'a LockedBumpAllocator {
    unsafe fn alloc(&mut self, layout: Layout) -> Result<*mut u8, AllocErr> {
        self.0.lock().alloc(layout)
    }
    unsafe fn dealloc(&mut self, ptr: *mut u8, layout: Layout) {
        self.0.lock().dealloc(ptr, layout)
    }
}

However, there is a more interesting solution for our bump allocator that avoids locking altogether. The idea is to exploit that we only need to update a single usize field byusing an AtomicUsize type. This type uses special synchronized hardware instructions to ensure data race freedom without requiring locks.

A lock-free Bump Allocator

A lock-free implementation looks like this:

use core::sync::atomic::{AtomicUsize, Ordering};
/// A simple allocator that allocates memory linearly and ignores freed memory.
#[derive(Debug)]
pub struct BumpAllocator {
    heap_start: usize,
    heap_end: usize,
    next: AtomicUsize,
}
impl BumpAllocator {
    pub const fn new(heap_start: usize, heap_end: usize) -> Self {
        // NOTE: requires adding #![feature(const_atomic_usize_new)] to lib.rs
        Self { heap_start, heap_end, next: AtomicUsize::new(heap_start) }
    }
}
unsafe impl<'a> Alloc for &'a BumpAllocator {
    unsafe fn alloc(&mut self, layout: Layout) -> Result<*mut u8, AllocErr> {
        loop {
            // load current state of the `next` field
            let current_next = self.next.load(Ordering::Relaxed);
            let alloc_start = align_up(current_next, layout.align());
            let alloc_end = alloc_start.saturating_add(layout.size());
            if alloc_end <= self.heap_end {
                // update the `next` pointer if it still has the value `current_next`
                let next_now = self.next.compare_and_swap(current_next, alloc_end,
                    Ordering::Relaxed);
                if next_now == current_next {
                    // next address was successfully updated, allocation succeeded
                    return Ok(alloc_start as *mut u8);
                }
            } else {
                return Err(AllocErr::Exhausted{ request: layout })
            }
        }
    }
    unsafe fn dealloc(&mut self, ptr: *mut u8, layout: Layout) {
        // do nothing, leak memory
    }
}

The implementation is a bit more complicated now. First, there is now a loop around the whole method body, since we might need multiple tries until we succeed (e.g. if multiple threads try to allocate at the same time). Also, the loads operation is an explicit method call now, i.e. self.next.load(Ordering::Relaxed) instead of just self.next. The ordering parameter makes it possible to restrict the automatic instruction reordering performed by both the compiler and the CPU itself. For example, it is used when implementing locks to ensure that no write to the locked variable happens before the lock is acquired. We don’t have such requirements, so we use the less restrictive Relaxed ordering.

The heart of this lock-free method is the compare_and_swap call that updates the next address:

...
let next_now = self.next.compare_and_swap(current_next, alloc_end,
    Ordering::Relaxed);
if next_now == current_next {
    // next address was successfully updated, allocation succeeded
    return Ok(alloc_start as *mut u8);
}
...

Compare-and-swap is a special CPU instruction that updates a variable with a given value if it still contains the value we expect. If it doesn’t, it means that another thread updated the value simultaneously, so we need to try again. The important feature is that this happens in a single uninteruptible operation (thus the name atomic), so no partial updates or intermediate states are possible.

In detail, compare_and_swap works by comparing next with the first argument and, in case they’re equal, updates next with the second parameter (the previous value is returned). To find out whether a switch happened, we check the returned previous value of next. If it is equal to the first parameter, the values were swapped. Otherwise, we try again in the next loop iteration.

Setting the Global Allocator

Now we can define a static bump allocator, that we can set as system allocator:

pub const HEAP_START: usize = 0o_000_001_000_000_0000;
pub const HEAP_SIZE: usize = 100 * 1024; // 100 KiB
#[global_allocator]
static HEAP_ALLOCATOR: BumpAllocator = BumpAllocator::new(HEAP_START,
    HEAP_START + HEAP_SIZE);

We use 0o_000_001_000_000_0000 as heap start address, which is the address starting at the second P3 entry. It doesn’t really matter which address we choose here as long as it’s unused. We use a heap size of 100 KiB, which should be large enough for the near future.

Putting the above in the memory::heap_allocator module would make most sense, but unfortunately there is currently a weird bug in the global allocator implementation that requires putting the global allocator in the root module. I hope it’s fixed soon, but until then we need to put the above lines in src/lib.rs. For that, we need to make the memory::heap_allocator module public and add an import for BumpAllocator. We also need to add the #![feature(global_allocator)] at the top of our lib.rs, since the global_allocator attribute is still unstable.

That’s it! We have successfully created and linked a custom system allocator. Now we’re ready to test it.

Testing

We should be able to allocate memory on the heap now. Let’s try it in our rust_main:

// in rust_main in src/lib.rs
use alloc::boxed::Box;
let heap_test = Box::new(42);

When we run it, a triple fault occurs and causes permanent rebooting. Let’s try debug it using QEMU and objdump as described in the previous post:

> qemu-system-x86_64 -d int -no-reboot -cdrom build/os-x86_64.iso
…
check_exception old: 0xffffffff new 0xe
     0: v=0e e=0002 i=0 cpl=0 IP=0008:0000000000102860 pc=0000000000102860
        SP=0010:0000000000116af0 CR2=0000000040000000
…

Aha! It’s a page fault (v=0e) and was caused by the code at 0x102860. The code tried to write (e=0002) to address 0x40000000. This address is 0o_000_001_000_000_0000 in octal, which is the HEAP_START address defined above. Of course it page-faults: We have forgotten to map the heap memory to some physical memory.

Some Refactoring

In order to map the heap cleanly, we do a bit of refactoring first. We move all memory initialization from our rust_main to a new memory::init function. Now our rust_main looks like this:

// in src/lib.rs
#[no_mangle]
pub extern "C" fn rust_main(multiboot_information_address: usize) {
    // ATTENTION: we have a very small stack and no guard page
    vga_buffer::clear_screen();
    println!("Hello World{}", "!");
    let boot_info = unsafe {
        multiboot2::load(multiboot_information_address)
    };
    enable_nxe_bit();
    enable_write_protect_bit();
    // set up guard page and map the heap pages
    memory::init(boot_info);
    use alloc::boxed::Box;
    let heap_test = Box::new(42);
    println!("It did not crash!");
    loop {}
}

The memory::init function looks like this:

// in src/memory/mod.rs
use multiboot2::BootInformation;
pub fn init(boot_info: &BootInformation) {
    let memory_map_tag = boot_info.memory_map_tag().expect(
        "Memory map tag required");
    let elf_sections_tag = boot_info.elf_sections_tag().expect(
        "Elf sections tag required");
    let kernel_start = elf_sections_tag.sections()
        .filter(|s| s.is_allocated()).map(|s| s.addr).min().unwrap();
    let kernel_end = elf_sections_tag.sections()
        .filter(|s| s.is_allocated()).map(|s| s.addr + s.size).max()
        .unwrap();
    println!("kernel start: {:#x}, kernel end: {:#x}",
             kernel_start,
             kernel_end);
    println!("multiboot start: {:#x}, multiboot end: {:#x}",
             boot_info.start_address(),
             boot_info.end_address());
    let mut frame_allocator = AreaFrameAllocator::new(
        kernel_start as usize, kernel_end as usize,
        boot_info.start_address(), boot_info.end_address(),
        memory_map_tag.memory_areas());
    paging::remap_the_kernel(&mut frame_allocator, boot_info);
}

We’ve just moved the code to a new function. However, we’ve sneaked some improvements in:

• An additional .filter(|s| s.is_allocated()) in the calculation of kernel_start and kernel_end. This ignores all sections that aren’t loaded to memory (such as debug sections). Thus, the kernel end address is no longer artificially increased by such sections.
• We use the start_address() and end_address() methods of boot_info instead of calculating the addresses manually.
• We use the alternate {:#x} form when printing kernel/multiboot addresses. Before, we used 0x{:x}, which leads to the same result. For a complete list of these “alternate” formatting forms, check out the std::fmt documentation.

Safety

It is important that the memory::init function is called only once, because it creates a new frame allocator based on kernel and multiboot start/end. When we call it a second time, a new frame allocator is created that reassigns the same frames, even if they are already in use.

In the second call it would use an identical frame allocator to remap the kernel. The remap_the_kernel function would request a frame from the frame allocator to create a new page table. But the returned frame is already in use, since we used it to create our current page table in the first call. In order to initialize the new table, the function zeroes it. This is the point where everything breaks, since we zero our current page table. The CPU is unable to read the next instruction and throws a page fault.

So we need to ensure that memory::init can be only called once. We could mark it as unsafe, which would bring it in line with Rust’s memory safety rules. However, that would just push the unsafety to the caller. The caller can still accidentally call the function twice, the only difference is that the mistake needs to happen inside unsafe blocks.

A better solution is to insert a check at the function’s beginning, that panics if the function is called a second time. This approach has a small runtime cost, but we only call it once, so it’s negligible. And we avoid two unsafe blocks (one at the calling site and one at the function itself), which is always good.

In order to make such checks easy, I created a small crate named once. To add it, we run cargo add once and add the following to our src/lib.rs:

// in src/lib.rs
#[macro_use]
extern crate once;

The crate provides an assert_has_not_been_called! macro (sorry for the long name :D). We can use it to fix the safety problem easily:

// in src/memory/mod.rs
pub fn init(boot_info: &BootInformation) {
    assert_has_not_been_called!("memory::init must be called only once");
    let memory_map_tag = ...
    ...
}

That’s it. Now our memory::init function can only be called once. The macro works by creating a static AtomicBool named CALLED, which is initialized to false. When the macro is invoked, it checks the value of CALLED and sets it to true. If the value was already true before, the macro panics.

Mapping the Heap

Now we’re ready to map the heap pages. In order to do it, we need access to the ActivePageTable or Mapper instance (see the page table and kernel remapping posts). For that we return it from the paging::remap_the_kernel function:

// in src/memory/paging/mod.rs
pub fn remap_the_kernel<A>(allocator: &mut A, boot_info: &BootInformation)
    -> ActivePageTable // new
    where A: FrameAllocator
{
    ...
    println!("guard page at {:#x}", old_p4_page.start_address());
    active_table // new
}

Now we have full page table access in the memory::init function. This allows us to map the heap pages to physical frames:

// in src/memory/mod.rs
pub fn init(boot_info: &BootInformation) {
    ...
    let mut frame_allocator = ...;
    // below is the new part
    let mut active_table = paging::remap_the_kernel(&mut frame_allocator,
        boot_info);
    use self::paging::Page;
    use {HEAP_START, HEAP_SIZE};
    let heap_start_page = Page::containing_address(HEAP_START);
    let heap_end_page = Page::containing_address(HEAP_START + HEAP_SIZE-1);
    for page in Page::range_inclusive(heap_start_page, heap_end_page) {
        active_table.map(page, paging::WRITABLE, &mut frame_allocator);
    }
}

The Page::range_inclusive function is just a copy of the Frame::range_inclusive function:

// in src/memory/paging/mod.rs
#[derive(…, PartialEq, Eq, PartialOrd, Ord)]
pub struct Page {...}
impl Page {
    ...
    pub fn range_inclusive(start: Page, end: Page) -> PageIter {
        PageIter {
            start: start,
            end: end,
        }
    }
}
pub struct PageIter {
    start: Page,
    end: Page,
}
impl Iterator for PageIter {
    type Item = Page;
    fn next(&mut self) -> Option<Page> {
        if self.start <= self.end {
            let page = self.start;
            self.start.number += 1;
            Some(page)
        } else {
            None
        }
    }
}

Now we map the whole heap to physical pages. This needs some time and might introduce a noticeable delay when we increase the heap size in the future. Another drawback is that we consume a large amount of physical frames even though we might not need the whole heap space. We will fix these problems in a future post by mapping the pages lazily.

It works!

Now Box and Vec should work. For example:

// in rust_main in src/lib.rs
use alloc::boxed::Box;
let mut heap_test = Box::new(42);
*heap_test -= 15;
let heap_test2 = Box::new("hello");
println!("{:?} {:?}", heap_test, heap_test2);
let mut vec_test = vec![1,2,3,4,5,6,7];
vec_test[3] = 42;
for i in &vec_test {
    print!("{} ", i);
}

We can also use all other types of the alloc crate, including:

• the reference counted pointers Rc and Arc
• the owned string type String and the format! macro
• Linked List
• the growable ring buffer VecDeque
• BinaryHeap
• BTreeMap and BTreeSet

A better Allocator

Right now, we leak every freed memory block. Thus, we run out of memory quickly, for example, by creating a new String in each iteration of a loop:

// in rust_main in src/lib.rs
for i in 0..10000 {
    format!("Some String");
}

To fix this, we need to create an allocator that keeps track of freed memory blocks and reuses them if possible. This introduces some challenges:

• We need to keep track of a possibly unlimited number of freed blocks. For example, an application could allocate n one-byte sized blocks and free every second block, which creates n/2 freed blocks. We can’t rely on any upper bound of freed block since n could be arbitrarily large.
• We can’t use any of the collections from above, since they rely on allocations themselves. (It might be possible as soon as RFC #1398 is implemented, which allows user-defined allocators for specific collection instances.)
• We need to merge adjacent freed blocks if possible. Otherwise, the freed memory is no longer usable for large allocations. We will discuss this point in more detail below.
• Our allocator should search the set of freed blocks quickly and keep fragmentation low.

Creating a List of freed Blocks

Where do we store the information about an unlimited number of freed blocks? We can’t use any fixed size data structure since it could always be too small for some allocation sequences. So we need some kind of dynamically growing set.

One possible solution could be to use an array-like data structure that starts at some unused virtual address. If the array becomes full, we increase its size and map new physical frames as backing storage. This approach would require a large part of the virtual address space since the array could grow significantly. We would need to create a custom implementation of a growable array and manipulate the page tables when deallocating. It would also consume a possibly large number of physical frames as backing storage.

We will choose another solution with different tradoffs. It’s not clearly “better” than the approach above and has significant disadvantages itself. However, it has one big advantage: It does not need any additional physical or virtual memory at all. This makes it less complex since we don’t need to manipulate any page tables. The idea is the following:

A freed memory block is not used anymore and no one needs the stored information. It is still mapped to a virtual address and backed by a physical page. So we just store the information about the freed block in the block itself. We keep a pointer to the first block and store a pointer to the next block in each block. Thus, we create a single linked list:

In the following, we call a freed block a hole. Each hole stores its size and a pointer to the next hole. If a hole is larger than needed, we leave the remaining memory unused. By storing a pointer to the first hole, we are able to traverse the complete list.

Initialization

When the heap is created, all of its memory is unused. Thus, it forms a single large hole:

The optional pointer to the next hole is set to None.

Allocation

In order to allocate a block of memory, we need to find a hole that satisfies the size and alignment requirements. If the found hole is larger than required, we split it into two smaller holes. For example, when we allocate a 24 byte block right after initialization, we split the single hole into a hole of size 24 and a hole with the remaining size:

Then we use the new 24 byte hole to perform the allocation:

To find a suitable hole, we can use several search strategies:

• best fit: Search the whole list and choose the smallest hole that satisfies the requirements.
• worst fit: Search the whole list and choose the largest hole that satisfies the requirements.
• first fit: Search the list from the beginning and choose the first hole that satisfies the requirements.

Each strategy has its advantages and disadvantages. Best fit uses the smallest hole possible and leaves larger holes for large allocations. But splitting the smallest hole might create a tiny hole, which is too small for most allocations. In contrast, the worst fit strategy always chooses the largest hole. Thus, it does not create tiny holes, but it consumes the large block, which might be required for large allocations.

For our use case, the best fit strategy is better than worst fit. The reason is that we have a minimal hole size of 16 bytes, since each hole needs to be able to store a size (8 bytes) and a pointer to the next hole (8 bytes). Thus, even the best fit strategy leads to holes of usable size. Furthermore, we will need to allocate very large blocks occasionally (e.g. for DMA buffers).

However, both best fit and worst fit have a significant problem: They need to scan the whole list for each allocation in order to find the optimal block. This leads to long allocation times if the list is long. The first fit strategy does not have this problem, as it returns as soon as it finds a suitable hole. It is fairly fast for small allocations and might only need to scan the whole list for large allocations.

Deallocation

To deallocate a block of memory, we can just insert its corresponding hole somewhere into the list. However, we need to merge adjacent holes. Otherwise, we are unable to reuse the freed memory for larger allocations. For example:

In order to use these adjacent holes for a large allocation, we need to merge them to a single large hole first:

The easiest way to ensure that adjacent holes are always merged, is to keep the hole list sorted by address. Thus, we only need to check the predecessor and the successor in the list when we free a memory block. If they are adjacent to the freed block, we merge the corresponding holes. Else, we insert the freed block as a new hole at the correct position.

Implementation

The detailed implementation would go beyond the scope of this post, since it contains several hidden difficulties. For example:

• Several merge cases: Merge with the previous hole, merge with the next hole, merge with both holes.
• We need to satisfy the alignment requirements, which requires additional splitting logic.
• The minimal hole size of 16 bytes: We must not create smaller holes when splitting a hole.

I created the linked_list_allocator crate to handle all of these cases. It consists of a Heap struct that provides an allocate_first_fit and a deallocate method. It also contains a LockedHeap type that wraps Heap into spinlock so that it’s usable as a static system allocator. If you are interested in the implementation details, check out the source code.

We need to add the extern crate to our Cargo.toml and our lib.rs:

> cargo add linked_list_allocator

// in src/lib.rs
extern crate linked_list_allocator;

Now we can change our global allocator:

use linked_list_allocator::LockedHeap;
#[global_allocator]
static HEAP_ALLOCATOR: LockedHeap = LockedHeap::empty();

We can’t initialize the linked list allocator statically, since it needs to initialize the first hole (like described above). This can’t be done at compile time, so the function can’t be a const function. Therefore we can only create an empty heap and initialize it later at runtime. For that, we add the following lines to our rust_main function:

// in src/lib.rs
#[no_mangle]
pub extern "C" fn rust_main(multiboot_information_address: usize) {
    […]
    // set up guard page and map the heap pages
    memory::init(boot_info);
    // initialize the heap allocator
    unsafe {
        HEAP_ALLOCATOR.lock().init(HEAP_START, HEAP_START + HEAP_SIZE);
    }
    […]
}

It is important that we initialize the heap after mapping the heap pages, since the init function writes to the heap memory (the first hole).

Our kernel uses the new allocator now, so we can deallocate memory without leaking it. The example from above should work now without causing an OOM situation:

// in rust_main in src/lib.rs
for i in 0..10000 {
    format!("Some String");
}

Performance

The linked list based approach has some performance problems. Each allocation or deallocation might need to scan the complete list of holes in the worst case. However, I think it’s good enough for now, since our heap will stay relatively small for the near future. When our allocator becomes a performance problem eventually, we can just replace it with a faster alternative.

Summary

Now we’re able to use heap storage in our kernel without leaking memory. This allows us to effectively process dynamic data such as user supplied strings in the future. We can also use Rc and Arc to create types with shared ownership. And we have access to various data structures such as Vec or Linked List, which will make our lives much easier. We even have some well tested and optimized binary heap and B-tree implementations!

What’s next?

This post concludes the section about memory management for now. We will revisit this topic eventually, but now it’s time to explore other topics. The upcoming posts will be about CPU exceptions and interrupts. We will catch all page, double, and triple faults and create a driver to read keyboard input. The next post starts by setting up a so-called Interrupt Descriptor Table.