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1001 lines
57 KiB
Markdown
1001 lines
57 KiB
Markdown
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title = "Paging Implementation"
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weight = 10
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path = "paging-implementation"
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date = 0000-01-01
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template = "second-edition/page.html"
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+++
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This post explains techniques to make the physical page table frames accessible to our kernel. It then uses such a technique to implement a function that translates virtual to physical addresses. It also explains how to create new mappings in the page tables.
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<!-- more -->
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This blog is openly developed on [GitHub]. If you have any problems or questions, please open an issue there. You can also leave comments [at the bottom]. The complete source code for this post can be found [here][post branch].
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[GitHub]: https://github.com/phil-opp/blog_os
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[at the bottom]: #comments
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[post branch]: https://github.com/phil-opp/blog_os/tree/post-10
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## Introduction
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In the [previous post] we learned about the principles of paging and how the 4-level page tables on x86_64 work. We also found out that the bootloader already set up a page table hierarchy for our kernel, which means that our kernel already runs on virtual addresses. This improves safety since illegal memory accesses cause page fault exceptions instead of modifying arbitrary physical memory.
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[previous post]: ./second-edition/posts/09-paging-introduction/index.md
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However, it also causes a problem when we try to access the page tables from our kernel because we can't directly access the physical addresses that are stored in page table entries or the `CR3` register. We experienced that problem already [at the end of the previous post] when we tried to inspect the active page tables.
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[at the end of the previous post]: ./second-edition/posts/09-paging-introduction/index.md#accessing-the-page-tables
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The next section discusses the problem in detail and provides different approaches to a solution. Afterward, we implement a function that traverses the page table hierarchy in order to translate virtual to physical addresses. Finally, we learn how to create new mappings in the page tables and how to find unused memory frames for creating new page tables.
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### Dependency Updates
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This post requires version 0.5.0 or later of the `x86_64` dependency and version 0.4.0 or later of the `bootloader` dependency. You can update the dependencies in your `Cargo.toml`:
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```toml
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[dependencies]
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x86_64 = "0.5.0"
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bootloader = "0.4.0"
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```
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## Accessing Page Tables
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Accessing the page tables from our kernel is not as easy as it may seem. To understand the problem let's take a look at the example 4-level page table hierarchy of the previous post again:
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The important thing here is that each page entry stores the _physical_ address of the next table. This avoids the need to run a translation for these addresses too, which would be bad for performance and could easily cause endless translation loops.
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The problem for us is that we can't directly access physical addresses from our kernel since our kernel also runs on top of virtual addresses. For example when we access address `4 KiB`, we access the _virtual_ address `4 KiB`, not the _physical_ address `4 KiB` where the level 4 page table is stored. When we want to access the physical address `4 KiB`, we can only do so through some virtual address that maps to it.
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So in order to access page table frames, we need to map some virtual pages to them. There are different ways to create these mappings that all allow us to access arbitrary page table frames:
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### Identity Mapping
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A simple solution is to **identity map all page tables**:
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In this example, we see various identity-mapped page table frames. This way the physical addresses of page tables are also valid virtual addresses so that we can easily access the page tables of all levels starting from the CR3 register.
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However, it clutters the virtual address space and makes it more difficult to find continuous memory regions of larger sizes. For example, imagine that we want to create a virtual memory region of size 1000 KiB in the above graphic, e.g. for [memory-mapping a file]. We can't start the region at `28 KiB` because it would collide with the already mapped page at `1004 MiB`. So we have to look further until we find a large enough unmapped area, for example at `1008 KiB`. This is a similar fragmentation problem as with [segmentation].
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[memory-mapping a file]: https://en.wikipedia.org/wiki/Memory-mapped_file
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[segmentation]: ./second-edition/posts/09-paging-introduction/index.md#fragmentation
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Equally, it makes it much more difficult to create new page tables, because we need to find physical frames whose corresponding pages aren't already in use. For example, let's assume that we reserved the _virtual_ 1000 KiB memory region starting at `1008 KiB` for our memory-mapped file. Now we can't use any frame with a _physical_ address between `1000 KiB` and `2008 KiB` anymore, because we can't identity map it.
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### Mapping at a Fixed Offset
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To avoid the problem of cluttering the virtual address space, we can **use a seperate memory region for page table mappings**. So instead of identity mapping page table frames, we map them at a fixed offset in the virtual address space. For example, the offset could be 10 TiB:
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By using the virtual memory in the range `10TiB..(10TiB + physical memory size)` exclusively for page table mappings, we avoid the collision problems of the identity mapping. Reserving such a large region of the virtual address space is only possible if the virtual address space is much larger than the physical memory size. This isn't a problem on x86_64 since the 48-bit address space is 256 TiB large.
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This approach still has the disadvantage that we need to create a new mapping whenever we create a new page table. Also, it does not allow accessing page tables of other address spaces, which would be useful when creating a new process.
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### Mapping the Complete Physical Memory
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We can solve these problems by **mapping the complete physical memory** instead of only page table frames:
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This approach allows our kernel to access arbitrary physical memory, including page table frames of other address spaces. The reserved virtual memory range has the same size as before, with the difference that it no longer contains unmapped pages.
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The disadvantage of this approach is that additional page tables are needed for storing the mapping of the physical memory. These page tables need to be stored somewhere, so they use up a part of physical memory, which can be a problem on devices with a small amount of memory.
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On x86_64, however, we can use [huge pages] with size 2MiB for the mapping, instead of the default 4KiB pages. This way, mapping 32 GiB of physical memory only requires 132 KiB for page tables since only one level 3 table and 32 level 2 tables are needed. Huge pages are also more cache efficient, since they use less entries in the translation lookaside buffer (TLB).
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[huge pages]: https://en.wikipedia.org/wiki/Page_%28computer_memory%29#Multiple_page_sizes
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### Temporary Mapping
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For devices with very small amounts of physical memory, we could **map the page tables frames only temporarily** when we need to access them. To be able to create the temporary mappings we only need a single identity-mapped level 1 table:
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The level 1 table in this graphic controls the first 2 MiB of the virtual address space. This is because it is reachable by starting at the CR3 register and following the 0th entry in the level 4, level 3, and level 2 page tables. The entry with index `8` maps the virtual page at address `32 KiB` to the physical frame at address `32 KiB`, thereby identity mapping the level 1 table itself. The graphic shows this identity-mapping by the horizontal arrow at `32 KiB`.
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By writing to the identity-mapped level 1 table, our kernel can create up to 511 temporary mappings (512 minus the entry required for the identity mapping). In the above example, the kernel mapped the 0th entry of the level 1 table to the frame with address `24 KiB`. This created a temporary mapping of the virtual page at `0 KiB` to the physical frame of the level 2 page table, indicated by the dashed arrow. Now the kernel can access the level 2 page table by writing to the page starting at `0 KiB`.
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The process for accessing an arbitrary page table frame with temporary mappings would be:
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- Search for a free entry in the identity-mapped level 1 table.
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- Map that entry to the physical frame of the page table that we want to access.
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- Access the target frame through the virtual page that maps to the entry.
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- Set the entry back to unused thereby removing the temporary mapping again.
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This approach reuses the same 512 virtual pages for creating the mappings and thus requires only 4KiB of physical memory. The drawback is that it is a bit cumbersome, especially since a new mapping might require modifications of multiple table levels, which means that we would need to repeat the above process multiple times.
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### Recursive Page Tables
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Another interesting approach, that requires no additional page tables at all, is to **map the page table recursively**. The idea behind this approach is to map some entry of the level 4 page table to the level 4 table itself. By doing this, we effectively reserve a part of the virtual address space and map all current and future page table frames to that space.
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Let's go through an example to understand how this all works:
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The only difference to the [example at the beginning of this post] is the additional entry at index `511` in the level 4 table, which is mapped to physical frame `4 KiB`, the frame of the level 4 table itself.
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[example at the beginning of this post]: #accessing-page-tables
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By letting the CPU follow this entry on a translation, it doesn't reach a level 3 table, but the same level 4 table again. This is similar to a recursive function that calls itself, therefore this table is called a _recursive page table_. The important thing is that the CPU assumes that every entry in the level 4 table points to a level 3 table, so it now treats the level 4 table as a level 3 table. This works because tables of all levels have the exact same layout on x86_64.
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By following the recursive entry one or multiple times before we start the actual translation, we can effectively shorten the number of levels that the CPU traverses. For example, if we follow the recursive entry once and then proceed to the level 3 table, the CPU thinks that the level 3 table is a level 2 table. Going further, it treats the level 2 table as a level 1 table and the level 1 table as the mapped frame. This means that we can now read and write the level 1 page table because the CPU thinks that it is the mapped frame. The graphic below illustrates the 5 translation steps:
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Similarly, we can follow the recursive entry twice before starting the translation to reduce the number of traversed levels to two:
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Let's go through it step by step: First, the CPU follows the recursive entry on the level 4 table and thinks that it reaches a level 3 table. Then it follows the recursive entry again and thinks that it reaches a level 2 table. But in reality, it is still on the level 4 table. When the CPU now follows a different entry, it lands on a level 3 table but thinks it is already on a level 1 table. So while the next entry points at a level 2 table, the CPU thinks that it points to the mapped frame, which allows us to read and write the level 2 table.
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Accessing the tables of levels 3 and 4 works in the same way. For accessing the level 3 table, we follow the recursive entry three times, tricking the CPU into thinking it is already on a level 1 table. Then we follow another entry and reach a level 3 table, which the CPU treats as a mapped frame. For accessing the level 4 table itself, we just follow the recursive entry four times until the CPU treats the level 4 table itself as mapped frame (in blue in the graphic below).
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It might take some time to wrap your head around the concept, but it works quite well in practice.
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In the section below we explain how to construct virtual addresses for following the recursive entry one or multiple times. We will not use recursive paging for our implementation, so you don't need to read it to continue with the post. If it interests you, just click on _"Address Calculation"_ to expand it.
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---
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<details>
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<summary><h4>Address Calculation</h4></summary>
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We saw that we can access tables of all levels by following the recursive entry once or multiple times before the actual translation. Since the indexes into the tables of the four levels are derived directly from the virtual address, we need to construct special virtual addresses for this technique. Remember, the page table indexes are derived from the address in the following way:
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Let's assume that we want to access the level 1 page table that maps a specific page. As we learned above, this means that we have to follow the recursive entry one time before continuing with the level 4, level 3, and level 2 indexes. To do that we move each block of the address one block to the right and set the original level 4 index to the index of the recursive entry:
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For accessing the level 2 table of that page, we move each index block two blocks to the right and set both the blocks of the original level 4 index and the original level 3 index to the index of the recursive entry:
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Accessing the level 3 table works by moving each block three blocks to the right and using the recursive index for the original level 4, level 3, and level 2 address blocks:
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Finally, we can access the level 4 table by moving each block four blocks to the right and using the recursive index for all address blocks except for the offset:
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We can now calculate virtual addresses for the page tables of all four levels. We can even calculate an address that points exactly to a specific page table entry by multiplying its index by 8, the size of a page table entry.
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The table below summarizes the address structure for accessing the different kinds of frames:
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Virtual Address for | Address Structure ([octal])
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------------------- | -------------------------------
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Page | `0o_SSSSSS_AAA_BBB_CCC_DDD_EEEE`
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Level 1 Table Entry | `0o_SSSSSS_RRR_AAA_BBB_CCC_DDDD`
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Level 2 Table Entry | `0o_SSSSSS_RRR_RRR_AAA_BBB_CCCC`
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Level 3 Table Entry | `0o_SSSSSS_RRR_RRR_RRR_AAA_BBBB`
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Level 4 Table Entry | `0o_SSSSSS_RRR_RRR_RRR_RRR_AAAA`
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[octal]: https://en.wikipedia.org/wiki/Octal
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Whereas `AAA` is the level 4 index, `BBB` the level 3 index, `CCC` the level 2 index, and `DDD` the level 1 index of the mapped frame, and `EEEE` the offset into it. `RRR` is the index of the recursive entry. When an index (three digits) is transformed to an offset (four digits), it is done by multiplying it by 8 (the size of a page table entry). With this offset, the resulting address directly points to the respective page table entry.
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`SSSSSS` are sign extension bits, which means that they are all copies of bit 47. This is a special requirement for valid addresses on the x86_64 architecture. We explained it in the [previous post][sign extension].
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[sign extension]: ./second-edition/posts/09-paging-introduction/index.md#paging-on-x86
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We use [octal] numbers for representing the addresses since each octal character represents three bits, which allows us to clearly separate the 9-bit indexes of the different page table levels. This isn't possible with the hexadecimal system where each character represents four bits.
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##### In Rust Code
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To construct such addresses in Rust code, you can use bitwise operations:
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```rust
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// the virtual address whose corresponding page tables you want to access
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let addr: usize = […];
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let r = 0o777; // recursive index
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let sign = 0o177777 << 48; // sign extension
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// retrieve the page table indices of the address that we want to translate
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let l4_idx = (addr >> 39) & 0o777; // level 4 index
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let l3_idx = (addr >> 30) & 0o777; // level 3 index
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let l2_idx = (addr >> 21) & 0o777; // level 2 index
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let l1_idx = (addr >> 12) & 0o777; // level 1 index
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let page_offset = addr & 0o7777;
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// calculate the table addresses
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let level_4_table_addr =
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sign | (r << 39) | (r << 30) | (r << 21) | (r << 12);
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let level_3_table_addr =
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sign | (r << 39) | (r << 30) | (r << 21) | (l4_idx << 12);
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let level_2_table_addr =
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sign | (r << 39) | (r << 30) | (l4_idx << 21) | (l3_idx << 12);
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let level_1_table_addr =
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sign | (r << 39) | (l4_idx << 30) | (l3_idx << 21) | (l2_idx << 12);
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```
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The above code assumes that the last level 4 entry with index `0o777` (511) is recursively mapped. This isn't the case currently, so the code won't work yet. See below on how to tell the bootloader to set up the recursive mapping.
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Alternatively to performing the bitwise operations by hand, you can use the [`RecursivePageTable`] type of the `x86_64` crate, which provides safe abstractions for various page table operations. For example, the code below shows how to translate a virtual address to its mapped physical address:
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[`RecursivePageTable`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.RecursivePageTable.html
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```rust
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// in src/memory.rs
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use x86_64::structures::paging::{Mapper, Page, PageTable, RecursivePageTable};
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use x86_64::{VirtAddr, PhysAddr};
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/// Creates a RecursivePageTable instance from the level 4 address.
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let level_4_table_addr = […];
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let level_4_table_ptr = level_4_table_addr as *mut PageTable;
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let recursive_page_table = unsafe {
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let level_4_table = &mut *level_4_table_ptr;
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RecursivePageTable::new(level_4_table).unwrap();
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}
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/// Retrieve the physical address for the given virtual address
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let addr: u64 = […]
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let addr = VirtAddr::new(addr);
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let page: Page = Page::containing_address(addr);
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// perform the translation
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let frame = recursive_page_table.translate_page(page);
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frame.map(|frame| frame.start_address() + u64::from(addr.page_offset()))
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```
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Again, a valid recursive mapping is required for this code. With such a mapping, the missing `level_4_table_addr` can be calculated as in the first code example.
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</details>
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---
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Recursive Paging is a interesting technique that shows how powerful a single mapping in a page table can be. It is relatively easy to implement and only requires a minimal amount of setup (just a single recursive entry), so it's a good choice for first experiments with paging.
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However, it also has some disadvantages:
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- It occupies a large amount of virtual memory (512GiB). This isn't a big problem in the large 48-bit address space, but it might lead to suboptimal cache behavior.
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- It only allows accessing the currently active address space easily. Accessing other address spaces is still possible by changing the recursive entry, but a temporary mapping is required for switching back. We described how to do this in the (outdated) [_Remap The Kernel_] post.
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- It heavily relies on the page table format of x86 and might not work on other architectures.
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[_Remap The Kernel_]: https://os.phil-opp.com/remap-the-kernel/#overview
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## Bootloader Support
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All of these approaches require page table modifications for their setup. For example, mappings for the physical memory need to be created or an entry of the level 4 table needs to be mapped recursively. The problem is that we can't create these required mappings without an existing way to access the page tables.
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This means that we need the help of the bootloader, which creates the page tables that our kernel runs on. The bootloader has access to the page tables, so it can create any mappings that we need. In its current implementation, the `bootloader` crate has support for two of the above approaches, controlled through [cargo features]:
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[cargo features]: https://doc.rust-lang.org/cargo/reference/manifest.html#the-features-section
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- With the `recursive-page-table` TODO feature, the bootloader maps an entry of the level 4 page table recursively. This allows the kernel to access the page tables as described in the [Recursive Page Tables](#recursive-page-tables) section.
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- The `map-physical-memory` TODO feature maps the complete physical memory somewhere into the virtual address space. Thus, the kernel can access all physical memory and can follow the [TODO](#todo) approach.
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We choose the TODO approach for our kernel since it is simple, platform-independent, and more powerful (it also allows to access non-page-table-frames). To enable the required bootloader support, we add the `map-physical-memory` feature to our `bootloader` dependency:
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```toml
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[dependencies]
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bootloader = { version = TODO, features = ["map-physical-memory"]}
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```
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With this feature enabled, the bootloader maps the complete physical memory to some unused virtual address range. To communicate the virtual address range to our kernel, the bootloader passes a _boot information_ structure.
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### Boot Information
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The `bootloader` crate defines a [`BootInfo`] struct that contains all the information it passes to our kernel. The struct is still in an early stage, so expect some breakage when updating to future [semver-incompatible] bootloader versions. With the `map-physical-memory` feature enabled, it currently has the two fields `memory_map` and `physical_memory_offset`:
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[`BootInfo`]: https://docs.rs/bootloader/0.3.11/bootloader/bootinfo/struct.BootInfo.html
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[semver-incompatible]: https://doc.rust-lang.org/stable/cargo/reference/specifying-dependencies.html#caret-requirements
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- The `memory_map` field contains an overview over the available physical memory. This tells our kernel how much physical memory is available in the system and which memory regions are reserved for devices such as the VGA hardware. The memory map can be queried from the BIOS or UEFI firmware, but only very early in the boot process. For this reason, it must be provided by the bootloader because there is no way for the kernel to retrieve it later. We will need the memory map later in this post.
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- The `physical_memory_offset` tells us the virtual start address of the physical memory mapping. By adding this offset to a physical address, we get the corresponding virtual address. This allows us to access arbitrary physical memory from our kernel.
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The bootloader passes the `BootInfo` struct to our kernel in form of a `&'static BootInfo` argument to our `_start` function. We don't have this argument declared in our function yet, so let's add it:
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```rust
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// in src/main.rs
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use bootloader::bootinfo::BootInfo;
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#[cfg(not(test))]
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#[no_mangle]
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pub extern "C" fn _start(boot_info: &'static BootInfo) -> ! { // new argument
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[…]
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}
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```
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It wasn't a problem to leave off this argument before because the x86_64 calling convention passes the first argument in a CPU register. Thus, the argument is simply ignored when it isn't declared. However, it would be a problem if we accidentally used a wrong argument type, since the compiler doesn't know the correct type signature of our entry point function.
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### The `entry_point` Macro
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Since our `_start` function is called externally from the bootloader, no checking of our function signature occurs. This means that we could let it take arbitrary arguments without any compilation errors, but it would fail or cause undefined behavior at runtime.
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To make sure that the entry point function has always the correct signature that the bootloader expects, the `bootloader` crate provides an [`entry_point`] macro that provides a type-checked way to define a Rust function as the entry point. Let's rewrite our entry point function to use this macro:
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[`entry_point`]: https://docs.rs/bootloader/0.3.12/bootloader/macro.entry_point.html
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```rust
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// in src/main.rs
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use bootloader::{bootinfo::BootInfo, entry_point};
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entry_point!(kernel_main);
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#[cfg(not(test))]
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fn kernel_main(boot_info: &'static BootInfo) -> ! {
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[…] // initialize GDT, IDT, PICS
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println!("It did not crash!");
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blog_os::hlt_loop();
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}
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```
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We no longer need to use `extern "C"` or `no_mangle` for our entry point, as the macro defines the real lower level `_start` entry point for us. The `kernel_main` function is now a completely normal Rust function, so we can choose an arbitrary name for it. The important thing is that it is type-checked so that a compilation error occurs when we now try to modify the function signature in any way, for example adding an argument or changing the argument type.
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## Implementation
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### Accessing the Page Tables
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### Translating Addresses
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||
|
||
### Creating a new Mapping
|
||
|
||
### Allocating Frames
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
---
|
||
|
||
|
||
The amount of physical memory and the memory regions reserved by devices like the VGA hardware vary between different machines. Only the BIOS or UEFI firmware knows exactly which memory regions can be used by the operating system and which regions are reserved. Both firmware standards provide functions to retrieve the memory map, but they can only be called very early in the boot process. For this reason, the bootloader already queries this and other information from the firmware.
|
||
|
||
|
||
---
|
||
|
||
|
||
|
||
|
||
For now it suffices to know that the bootloader used a technique called _recursive page tables_ to map the last page of the virtual address space to the physical frame of the level 4 page table. The last page of the virtual address space is `0xffff_ffff_ffff_f000`, so let's use it to read some entries of that table:
|
||
|
||
```rust
|
||
// in src/main.rs
|
||
|
||
#[cfg(not(test))]
|
||
#[no_mangle]
|
||
pub extern "C" fn _start() -> ! {
|
||
[…] // initialize GDT, IDT, PICS
|
||
|
||
let level_4_table_pointer = 0xffff_ffff_ffff_f000 as *const u64;
|
||
for i in 0..10 {
|
||
let entry = unsafe { *level_4_table_pointer.offset(i) };
|
||
println!("Entry {}: {:#x}", i, entry);
|
||
}
|
||
|
||
println!("It did not crash!");
|
||
blog_os::hlt_loop();
|
||
}
|
||
```
|
||
|
||
We cast the address of the last virtual page to a pointer to an `u64`. As we saw in the [previous section][page table format], each page table entry is 8 bytes (64 bits), so an `u64` represents exactly one entry. We print the first 10 entries of the table using a `for` loop. Inside the loop, we use an unsafe block to read from the raw pointer and the [`offset` method] to perform pointer arithmetic.
|
||
|
||
[page table format]: #page-table-format
|
||
[`offset` method]: https://doc.rust-lang.org/std/primitive.pointer.html#method.offset
|
||
|
||
When we run it, we see the following output:
|
||
|
||
|
||
|
||
When we look at the [format of page table entries][page table format], we see that the value `0x2023` of entry 0 means that the entry is `present`, `writable`, was `accessed` by the CPU, and is mapped to frame `0x2000`. Entry 1 is mapped to frame `0x6e2000` and has the same flags as entry 0, with the addition of the `dirty` flag that indicates that the page was written. Entries 2–9 are not `present`, so these virtual address ranges are not mapped to any physical addresses.
|
||
|
||
Instead of working with unsafe raw pointers we can use the [`PageTable`] type of the `x86_64` crate:
|
||
|
||
[`PageTable`]: https://docs.rs/x86_64/0.3.4/x86_64/structures/paging/struct.PageTable.html
|
||
|
||
```rust
|
||
// in src/main.rs
|
||
|
||
#[cfg(not(test))]
|
||
#[no_mangle]
|
||
pub extern "C" fn _start() -> ! {
|
||
[…] // initialize GDT, IDT, PICS
|
||
|
||
use x86_64::structures::paging::PageTable;
|
||
|
||
let level_4_table_ptr = 0xffff_ffff_ffff_f000 as *const PageTable;
|
||
let level_4_table = unsafe {&*level_4_table_ptr};
|
||
for i in 0..10 {
|
||
println!("Entry {}: {:?}", i, level_4_table[i]);
|
||
}
|
||
|
||
println!("It did not crash!");
|
||
blog_os::hlt_loop();
|
||
}
|
||
```
|
||
|
||
Here we cast the `0xffff_ffff_ffff_f000` pointer first to a raw pointer and then transform it to a Rust reference. This operation still needs `unsafe`, because the compiler can't know that this accessing this address is valid. But after this conversion we have a safe `&PageTable` type, which allows us to access the individual entries through safe, bounds checked [indexing operations].
|
||
|
||
[indexing operations]: https://doc.rust-lang.org/core/ops/trait.Index.html
|
||
|
||
The crate also provides some abstractions for the individual entries so that we directly see which flags are set when we print them:
|
||
|
||
|
||
|
||
The next step would be to follow the pointers in entry 0 or entry 1 to a level 3 page table. But we now again have the problem that `0x2000` and `0x6e5000` are physical addresses, so we can't access them directly. This problem will be solved in the next post.
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
## Implementation
|
||
|
||
After all this theory we can finally start our implementation. Conveniently, the bootloader not only created page tables for our kernel, but it also created a recursive mapping in the last entry of the level 4 table. The bootloader did this because otherwise there would be a [chicken or egg problem]: We need to access the level 4 table to create a recursive mapping, but we can't access it without some kind of mapping.
|
||
|
||
[chicken or egg problem]: https://en.wikipedia.org/wiki/Chicken_or_the_egg
|
||
|
||
We already used this recursive mapping [at the end of the previous post] to access the level 4 table. We did this through the hardcoded address `0xffff_ffff_ffff_f000`. When we convert this address to [octal] and compare it with the above table, we can see that it exactly follows the structure of a level 4 table entry with `RRR` = `0o777`, `AAAA` = 0, and the sign extension bits set to `1` each:
|
||
|
||
```
|
||
structure: 0o_SSSSSS_RRR_RRR_RRR_RRR_AAAA
|
||
address: 0o_177777_777_777_777_777_0000
|
||
```
|
||
|
||
With our knowledge about recursive page tables we can now create virtual addresses to access all active page tables. This allows us to create a translation function in software.
|
||
|
||
### Translating Addresses
|
||
|
||
As a first step, let's create a function that translates a virtual address to a physical address by walking the page table hierarchy:
|
||
|
||
```rust
|
||
// in src/lib.rs
|
||
|
||
pub mod memory;
|
||
```
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
use x86_64::PhysAddr;
|
||
use x86_64::structures::paging::PageTable;
|
||
|
||
/// Returns the physical address for the given virtual address, or `None` if the
|
||
/// virtual address is not mapped.
|
||
pub fn translate_addr(addr: usize) -> Option<PhysAddr> {
|
||
// introduce variables for the recursive index and the sign extension bits
|
||
// TODO: Don't hardcode these values
|
||
let r = 0o777; // recursive index
|
||
let sign = 0o177777 << 48; // sign extension
|
||
|
||
// retrieve the page table indices of the address that we want to translate
|
||
let l4_idx = (addr >> 39) & 0o777; // level 4 index
|
||
let l3_idx = (addr >> 30) & 0o777; // level 3 index
|
||
let l2_idx = (addr >> 21) & 0o777; // level 2 index
|
||
let l1_idx = (addr >> 12) & 0o777; // level 1 index
|
||
let page_offset = addr & 0o7777;
|
||
|
||
// calculate the table addresses
|
||
let level_4_table_addr =
|
||
sign | (r << 39) | (r << 30) | (r << 21) | (r << 12);
|
||
let level_3_table_addr =
|
||
sign | (r << 39) | (r << 30) | (r << 21) | (l4_idx << 12);
|
||
let level_2_table_addr =
|
||
sign | (r << 39) | (r << 30) | (l4_idx << 21) | (l3_idx << 12);
|
||
let level_1_table_addr =
|
||
sign | (r << 39) | (l4_idx << 30) | (l3_idx << 21) | (l2_idx << 12);
|
||
|
||
// check that level 4 entry is mapped
|
||
let level_4_table = unsafe { &*(level_4_table_addr as *const PageTable) };
|
||
if level_4_table[l4_idx].addr().is_null() {
|
||
return None;
|
||
}
|
||
|
||
// check that level 3 entry is mapped
|
||
let level_3_table = unsafe { &*(level_3_table_addr as *const PageTable) };
|
||
if level_3_table[l3_idx].addr().is_null() {
|
||
return None;
|
||
}
|
||
|
||
// check that level 2 entry is mapped
|
||
let level_2_table = unsafe { &*(level_2_table_addr as *const PageTable) };
|
||
if level_2_table[l2_idx].addr().is_null() {
|
||
return None;
|
||
}
|
||
|
||
// check that level 1 entry is mapped and retrieve physical address from it
|
||
let level_1_table = unsafe { &*(level_1_table_addr as *const PageTable) };
|
||
let phys_addr = level_1_table[l1_idx].addr();
|
||
if phys_addr.is_null() {
|
||
return None;
|
||
}
|
||
|
||
Some(phys_addr + page_offset)
|
||
}
|
||
```
|
||
|
||
First, we introduce variables for the recursive index (511 = `0o777`) and the sign extension bits (which are 1 each). Then we calculate the page table indices and the page offset from the address through bitwise operations as specified in the graphic:
|
||
|
||
|
||
|
||
In the next step we calculate the virtual addresses of the four page tables as descripbed in the [address calculation] section. We transform each of these addresses to [`PageTable`] references later in the function. These transformations are `unsafe` operations since the compiler can't know that these addresses are valid.
|
||
|
||
[address calculation]: #address-calculation
|
||
[`PageTable`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.PageTable.html
|
||
|
||
After the address calculation, we use the indexing operator to look at the entry in the level 4 table. If that entry is null, there is no level 3 table for this level 4 entry, which means that the `addr` is not mapped to any physical memory, so we return `None`. If the entry is not `None`, we know that a level 3 table exists. We then do the same cast and entry-checking as with the level 4 table.
|
||
|
||
After we checked the three higher level pages, we can finally read the entry of the level 1 table that tells us the physical frame that the address is mapped to. As the last step, we add the page offset to that address and return it.
|
||
|
||
If we knew that the address is mapped, we could directly access the level 1 table without looking at the higher level pages first. But since we don't know this, we have to check whether the level 1 table exists first, otherwise our function would cause a page fault for unmapped addresses.
|
||
|
||
#### Try it out
|
||
|
||
We can use our new translation function to translate some virtual addresses in our `_start` function:
|
||
|
||
```rust
|
||
// in src/main.rs
|
||
|
||
#[cfg(not(test))]
|
||
#[no_mangle]
|
||
pub extern "C" fn _start() -> ! {
|
||
[…] // initialize GDT, IDT, PICS
|
||
|
||
use blog_os::memory::translate_addr;
|
||
|
||
// the identity-mapped vga buffer page
|
||
println!("0xb8000 -> {:?}", translate_addr(0xb8000));
|
||
// some code page
|
||
println!("0x20010a -> {:?}", translate_addr(0x20010a));
|
||
// some stack page
|
||
println!("0x57ac001ffe48 -> {:?}", translate_addr(0x57ac001ffe48));
|
||
|
||
|
||
println!("It did not crash!");
|
||
blog_os::hlt_loop();
|
||
}
|
||
```
|
||
|
||
When we run it, we see the following output:
|
||
|
||
|
||
|
||
As expected, the identity-mapped address `0xb8000` translates to the same physical address. The code page and the stack page translate to some arbitrary physical addresses, which depend on how the bootloader created the initial mapping for our kernel.
|
||
|
||
#### The `RecursivePageTable` Type
|
||
|
||
The `x86_64` provides a [`RecursivePageTable`] type that implements safe abstractions for various page table operations. By using this type, we can reimplement our `translate_addr` function in a much cleaner way:
|
||
|
||
[`RecursivePageTable`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.RecursivePageTable.html
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
use x86_64::structures::paging::{Mapper, Page, PageTable, RecursivePageTable};
|
||
use x86_64::{VirtAddr, PhysAddr};
|
||
|
||
/// Creates a RecursivePageTable instance from the level 4 address.
|
||
///
|
||
/// This function is unsafe because it can break memory safety if an invalid
|
||
/// address is passed.
|
||
pub unsafe fn init(level_4_table_addr: usize) -> RecursivePageTable<'static> {
|
||
let level_4_table_ptr = level_4_table_addr as *mut PageTable;
|
||
let level_4_table = &mut *level_4_table_ptr;
|
||
RecursivePageTable::new(level_4_table).unwrap()
|
||
}
|
||
|
||
/// Returns the physical address for the given virtual address, or `None` if
|
||
/// the virtual address is not mapped.
|
||
pub fn translate_addr(addr: u64, recursive_page_table: &RecursivePageTable)
|
||
-> Option<PhysAddr>
|
||
{
|
||
let addr = VirtAddr::new(addr);
|
||
let page: Page = Page::containing_address(addr);
|
||
|
||
// perform the translation
|
||
let frame = recursive_page_table.translate_page(page);
|
||
frame.map(|frame| frame.start_address() + u64::from(addr.page_offset()))
|
||
}
|
||
```
|
||
|
||
The `RecursivePageTable` type encapsulates the unsafety of the page table walk completely so that we no longer need `unsafe` in our `translate_addr` function. The `init` function needs to be unsafe because the caller has to guarantee that the passed `level_4_table_addr` is valid.
|
||
|
||
Our `_start` function needs to be updated for the new function signature in the following way:
|
||
|
||
```rust
|
||
// in src/main.rs
|
||
|
||
#[cfg(not(test))]
|
||
#[no_mangle]
|
||
pub extern "C" fn _start() -> ! {
|
||
[…] // initialize GDT, IDT, PICS
|
||
|
||
use blog_os::memory::{self, translate_addr};
|
||
|
||
const LEVEL_4_TABLE_ADDR: usize = 0o_177777_777_777_777_777_0000;
|
||
let recursive_page_table = unsafe { memory::init(LEVEL_4_TABLE_ADDR) };
|
||
|
||
// the identity-mapped vga buffer page
|
||
println!("0xb8000 -> {:?}", translate_addr(0xb8000, &recursive_page_table));
|
||
// some code page
|
||
println!("0x20010a -> {:?}", translate_addr(0x20010a, &recursive_page_table));
|
||
// some stack page
|
||
println!("0x57ac001ffe48 -> {:?}", translate_addr(0x57ac001ffe48,
|
||
&recursive_page_table));
|
||
|
||
println!("It did not crash!");
|
||
blog_os::hlt_loop();
|
||
}
|
||
```
|
||
|
||
Instead of passing the `LEVEL_4_TABLE_ADDR` to `translate_addr` and accessing the page tables through unsafe raw pointers, we now pass references to a `RecursivePageTable` type. By doing this, we now have a safe abstraction and clear ownership semantics. This ensures that we can't accidentally modify the page table concurrently, because an exclusive borrow of the `RecursivePageTable` is needed in order to modify it.
|
||
|
||
When we run it, we see the same result as with our handcrafted translation function.
|
||
|
||
#### Making Unsafe Functions Safer
|
||
|
||
Our `memory::init` function is an [unsafe function], which means that an `unsafe` block is required for calling it because the caller has to guarantee that certain requirements are met. In our case, the requirement is that the passed address is mapped to the physical frame of the level 4 page table.
|
||
|
||
[unsafe function]: https://doc.rust-lang.org/book/ch19-01-unsafe-rust.html#calling-an-unsafe-function-or-method
|
||
|
||
The second property of unsafe functions is that their complete body is treated as an `unsafe` block, which means that it can perform all kinds of unsafe operations without additional unsafe blocks. This is the reason that we didn't need an `unsafe` block for dereferencing the raw `level_4_table_ptr`:
|
||
|
||
```rust
|
||
pub unsafe fn init(level_4_table_addr: usize) -> RecursivePageTable<'static> {
|
||
let level_4_table_ptr = level_4_table_addr as *mut PageTable;
|
||
let level_4_table = &mut *level_4_table_ptr; // <- this operation is unsafe
|
||
RecursivePageTable::new(level_4_table).unwrap()
|
||
}
|
||
```
|
||
|
||
The problem with this is that we don't immediately see which parts are unsafe. For example, we don't know whether the `RecursivePageTable::new` function is unsafe or not without looking at [its definition][RecursivePageTable::new]. This makes it very easy to accidentally do something unsafe without noticing.
|
||
|
||
[RecursivePageTable::new]: https://docs.rs/x86_64/0.3.6/x86_64/structures/paging/struct.RecursivePageTable.html#method.new
|
||
|
||
To avoid this problem, we can add a safe inner function:
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
pub unsafe fn init(level_4_table_addr: usize) -> RecursivePageTable<'static> {
|
||
/// Rust currently treats the whole body of unsafe functions as an unsafe
|
||
/// block, which makes it difficult to see which operations are unsafe. To
|
||
/// limit the scope of unsafe we use a safe inner function.
|
||
fn init_inner(level_4_table_addr: usize) -> RecursivePageTable<'static> {
|
||
let level_4_table_ptr = level_4_table_addr as *mut PageTable;
|
||
let level_4_table = unsafe { &mut *level_4_table_ptr };
|
||
RecursivePageTable::new(level_4_table).unwrap()
|
||
}
|
||
|
||
init_inner(level_4_table_addr)
|
||
}
|
||
```
|
||
|
||
Now an `unsafe` block is required again for dereferencing the `level_4_table_ptr` and we immediately see that this is the only unsafe operations in the function. There is currently an open [RFC][unsafe-fn-rfc] to change this unfortunate property of unsafe functions that would allow us to avoid the above boilerplate.
|
||
|
||
[unsafe-fn-rfc]: https://github.com/rust-lang/rfcs/pull/2585
|
||
|
||
### Creating a new Mapping
|
||
After reading the page tables and creating a translation function, the next step is to create a new mapping in the page table hierarchy.
|
||
|
||
The difficulty of creating a new mapping depends on the virtual page that we want to map. In the easiest case, the level 1 page table for the page already exists and we just need to write a single entry. In the most difficult case, the page is in a memory region for that no level 3 exists yet so that we need to create new level 3, level 2 and level 1 page tables first.
|
||
|
||
Let's start with the simple case and assume that we don't need to create new page tables. The bootloader loads itself in the first megabyte of the virtual address space, so we know that a valid level 1 table exists for this region. We can choose any unused page in this memory region for our example mapping, for example, the page at address `0x1000`. As the target frame we use `0xb8000`, the frame of the VGA text buffer. This way we can easily test whether our mapping worked.
|
||
|
||
We implement it in a new `create_mapping` function in our `memory` module:
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
use x86_64::structures::paging::{FrameAllocator, PhysFrame, Size4KiB};
|
||
|
||
pub fn create_example_mapping(
|
||
recursive_page_table: &mut RecursivePageTable,
|
||
frame_allocator: &mut impl FrameAllocator<Size4KiB>,
|
||
) {
|
||
use x86_64::structures::paging::PageTableFlags as Flags;
|
||
|
||
let page: Page = Page::containing_address(VirtAddr::new(0x1000));
|
||
let frame = PhysFrame::containing_address(PhysAddr::new(0xb8000));
|
||
let flags = Flags::PRESENT | Flags::WRITABLE;
|
||
|
||
let map_to_result = unsafe {
|
||
recursive_page_table.map_to(page, frame, flags, frame_allocator)
|
||
};
|
||
map_to_result.expect("map_to failed").flush();
|
||
}
|
||
```
|
||
|
||
The function takes a mutable reference to the `RecursivePageTable` because it needs to modify it and a `FrameAllocator` that is explained below. It then uses the [`map_to`] function of the [`Mapper`] trait to map the page at address `0x1000` to the physical frame at address `0xb8000`. The function is unsafe because it's possible to break memory safety with invalid arguments.
|
||
|
||
[`map_to`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/trait.Mapper.html#tymethod.map_to
|
||
[`Mapper`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/trait.Mapper.html
|
||
|
||
Apart from the `page` and `frame` arguments, the [`map_to`] function takes two more arguments. The third argument is a set of flags for the page table entry. We set the `PRESENT` flag because it is required for all valid entries and the `WRITABLE` flag to make the mapped page writable.
|
||
|
||
The fourth argument needs to be some structure that implements the [`FrameAllocator`] trait. The `map_to` method needs this argument because it might need unused frames for creating new page tables. The `Size4KiB` argument in the trait implementation is needed because the [`Page`] and [`PhysFrame`] types are [generic] over the [`PageSize`] trait to work with both standard 4KiB pages and huge 2MiB/1GiB pages.
|
||
|
||
[`FrameAllocator`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/trait.FrameAllocator.html
|
||
[`Page`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.Page.html
|
||
[`PhysFrame`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.PhysFrame.html
|
||
[generic]: https://doc.rust-lang.org/book/ch10-00-generics.html
|
||
[`PageSize`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/trait.PageSize.html
|
||
|
||
The `map_to` function can fail, so it returns a [`Result`]. Since this is just some example code that does not need to be robust, we just use [`expect`] to panic when an error occurs. On success, the function returns a [`MapperFlush`] type that provides an easy way to flush the newly mapped page from the translation lookaside buffer (TLB) with its [`flush`] method. Like `Result`, the type uses the [`#[must_use]`] attribute to emit a warning when we accidentally forget to use it.
|
||
|
||
[`Result`]: https://doc.rust-lang.org/core/result/enum.Result.html
|
||
[`expect`]: https://doc.rust-lang.org/core/result/enum.Result.html#method.expect
|
||
[`MapperFlush`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.MapperFlush.html
|
||
[`flush`]: https://docs.rs/x86_64/0.3.5/x86_64/structures/paging/struct.MapperFlush.html#method.flush
|
||
[`#[must_use]`]: https://doc.rust-lang.org/std/result/#results-must-be-used
|
||
|
||
Since we know that no new page tables are required for the address `0x1000`, a frame allocator that always returns `None` suffices. We create such an `EmptyFrameAllocator` for testing our mapping function:
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
/// A FrameAllocator that always returns `None`.
|
||
pub struct EmptyFrameAllocator;
|
||
|
||
impl FrameAllocator<Size4KiB> for EmptyFrameAllocator {
|
||
fn allocate_frame(&mut self) -> Option<PhysFrame> {
|
||
None
|
||
}
|
||
}
|
||
```
|
||
|
||
(If you're getting a 'method `allocate_frame` is not a member of trait `FrameAllocator`' error, you need to update `x86_64` to version 0.4.0.)
|
||
|
||
We can now test the new mapping function in our `main.rs`:
|
||
|
||
```rust
|
||
// in src/main.rs
|
||
|
||
#[cfg(not(test))]
|
||
#[no_mangle]
|
||
pub extern "C" fn _start() -> ! {
|
||
[…] // initialize GDT, IDT, PICS
|
||
|
||
use blog_os::memory::{create_example_mapping, EmptyFrameAllocator};
|
||
|
||
const LEVEL_4_TABLE_ADDR: usize = 0o_177777_777_777_777_777_0000;
|
||
let mut recursive_page_table = unsafe { memory::init(LEVEL_4_TABLE_ADDR) };
|
||
|
||
create_example_mapping(&mut recursive_page_table, &mut EmptyFrameAllocator);
|
||
unsafe { (0x1900 as *mut u64).write_volatile(0xf021f077f065f04e)};
|
||
|
||
println!("It did not crash!");
|
||
blog_os::hlt_loop();
|
||
}
|
||
```
|
||
|
||
We first create the mapping for the page at `0x1000` by calling our `create_example_mapping` function with a mutable reference to the `RecursivePageTable` instance. This maps the page `0x1000` to the VGA text buffer, so we should see any write to it on the screen.
|
||
|
||
Then we write the value `0xf021f077f065f04e` to this page, which represents the string _"New!"_ on white background. We don't write directly to the beginning of the page at `0x1000` since the top line is directly shifted off the screen by the next `println`. Instead, we write to offset `0x900`, which is about in the middle of the screen. As we learned [in the _“VGA Text Mode”_ post], writes to the VGA buffer should be volatile, so we use the [`write_volatile`] method.
|
||
|
||
[in the _“VGA Text Mode”_ post]: ./second-edition/posts/03-vga-text-buffer/index.md#volatile
|
||
[`write_volatile`]: https://doc.rust-lang.org/std/primitive.pointer.html#method.write_volatile
|
||
|
||
When we run it in QEMU, we see the following output:
|
||
|
||
|
||
|
||
The _"New!"_ on the screen is by our write to `0x1900`, which means that we successfully created a new mapping in the page tables.
|
||
|
||
This only worked because there was already a level 1 table for mapping page `0x1000`. When we try to map a page for that no level 1 table exists yet, the `map_to` function fails because it tries to allocate frames from the `EmptyFrameAllocator` for creating new page tables. We can see that happen when we try to map page `0xdeadbeaf000` instead of `0x1000`:
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
pub fn create_example_mapping(…) {
|
||
[…]
|
||
let page: Page = Page::containing_address(VirtAddr::new(0xdeadbeaf000));
|
||
[…]
|
||
}
|
||
|
||
// in src/main.rs
|
||
|
||
#[no_mangle]
|
||
pub extern "C" fn _start() -> ! {
|
||
[…]
|
||
unsafe { (0xdeadbeaf900 as *mut u64).write_volatile(0xf021f077f065f04e)};
|
||
[…]
|
||
}
|
||
```
|
||
|
||
When we run it, a panic with the following error message occurs:
|
||
|
||
```
|
||
panicked at 'map_to failed: FrameAllocationFailed', /…/result.rs:999:5
|
||
```
|
||
|
||
To map pages that don't have a level 1 page table yet we need to create a proper `FrameAllocator`. But how do we know which frames are unused and how much physical memory is available?
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
---
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
|
||
### Allocating Frames
|
||
|
||
Now that we have access to the memory map through the boot information we can create a proper frame allocator on top. We start with a generic skeleton:
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
pub struct BootInfoFrameAllocator<I> where I: Iterator<Item = PhysFrame> {
|
||
frames: I,
|
||
}
|
||
|
||
impl<I> FrameAllocator<Size4KiB> for BootInfoFrameAllocator<I>
|
||
where I: Iterator<Item = PhysFrame>
|
||
{
|
||
fn allocate_frame(&mut self) -> Option<PhysFrame> {
|
||
self.frames.next()
|
||
}
|
||
}
|
||
```
|
||
|
||
The `frames` field can be initialized with an arbitrary [`Iterator`] of frames. This allows us to just delegate `alloc` calls to the [`Iterator::next`] method.
|
||
|
||
[`Iterator`]: https://doc.rust-lang.org/core/iter/trait.Iterator.html
|
||
[`Iterator::next`]: https://doc.rust-lang.org/core/iter/trait.Iterator.html#tymethod.next
|
||
|
||
The initialization of the `BootInfoFrameAllocator` happens in a new `init_frame_allocator` function:
|
||
|
||
```rust
|
||
// in src/memory.rs
|
||
|
||
use bootloader::bootinfo::{MemoryMap, MemoryRegionType};
|
||
|
||
/// Create a FrameAllocator from the passed memory map
|
||
pub fn init_frame_allocator(
|
||
memory_map: &'static MemoryMap,
|
||
) -> BootInfoFrameAllocator<impl Iterator<Item = PhysFrame>> {
|
||
// get usable regions from memory map
|
||
let regions = memory_map
|
||
.iter()
|
||
.filter(|r| r.region_type == MemoryRegionType::Usable);
|
||
// map each region to its address range
|
||
let addr_ranges = regions.map(|r| r.range.start_addr()..r.range.end_addr());
|
||
// transform to an iterator of frame start addresses
|
||
let frame_addresses = addr_ranges.flat_map(|r| r.into_iter().step_by(4096));
|
||
// create `PhysFrame` types from the start addresses
|
||
let frames = frame_addresses.map(|addr| {
|
||
PhysFrame::containing_address(PhysAddr::new(addr))
|
||
});
|
||
|
||
BootInfoFrameAllocator { frames }
|
||
}
|
||
```
|
||
|
||
This function uses iterator combinator methods to transform the initial `MemoryMap` into an iterator of usable physical frames:
|
||
|
||
- First, we call the `iter` method to convert the memory map to an iterator of [`MemoryRegion`]s. Then we use the [`filter`] method to skip any reserved or otherwise unavailable regions. The bootloader updates the memory map for all the mappings it creates, so frames that are used by our kernel (code, data or stack) or to store the boot information are already marked as `InUse` or similar. Thus we can be sure that `Usable` frames are not used somewhere else.
|
||
- In the second step, we use the [`map`] combinator and Rust's [range syntax] to transform our iterator of memory regions to an iterator of address ranges.
|
||
- The third step is the most complicated: We convert each range to an iterator through the `into_iter` method and then choose every 4096th address using [`step_by`]. Since 4096 bytes (= 4 KiB) is the page size, we get the start address of each frame. The bootloader page aligns all usable memory areas so that we don't need any alignment or rounding code here. By using [`flat_map`] instead of `map`, we get an `Iterator<Item = u64>` instead of an `Iterator<Item = Iterator<Item = u64>>`.
|
||
- In the final step, we convert the start addresses to `PhysFrame` types to construct the desired `Iterator<Item = PhysFrame>`. We then use this iterator to create and return a new `BootInfoFrameAllocator`.
|
||
|
||
[`MemoryRegion`]: https://docs.rs/bootloader/0.3.12/bootloader/bootinfo/struct.MemoryRegion.html
|
||
[`filter`]: https://doc.rust-lang.org/core/iter/trait.Iterator.html#method.filter
|
||
[`map`]: https://doc.rust-lang.org/core/iter/trait.Iterator.html#method.map
|
||
[range syntax]: https://doc.rust-lang.org/core/ops/struct.Range.html
|
||
[`step_by`]: https://doc.rust-lang.org/core/iter/trait.Iterator.html#method.step_by
|
||
[`flat_map`]: https://doc.rust-lang.org/core/iter/trait.Iterator.html#method.flat_map
|
||
|
||
We can now modify our `kernel_main` function to pass a `BootInfoFrameAllocator` instance instead of an `EmptyFrameAllocator`:
|
||
|
||
```rust
|
||
// in src/main.rs
|
||
|
||
#[cfg(not(test))]
|
||
fn kernel_main(boot_info: &'static BootInfo) -> ! {
|
||
[…] // initialize GDT, IDT, PICS
|
||
|
||
use x86_64::structures::paging::{PageTable, RecursivePageTable};
|
||
|
||
let mut recursive_page_table = unsafe {
|
||
memory::init(boot_info.p4_table_addr as usize)
|
||
};
|
||
// new
|
||
let mut frame_allocator = memory::init_frame_allocator(&boot_info.memory_map);
|
||
|
||
blog_os::memory::create_mapping(&mut recursive_page_table, &mut frame_allocator);
|
||
unsafe { (0xdeadbeaf900 as *mut u64).write_volatile(0xf021f077f065f04e)};
|
||
|
||
println!("It did not crash!");
|
||
blog_os::hlt_loop();
|
||
}
|
||
```
|
||
|
||
Now the mapping succeeds and we see the black-on-white _"New!"_ on the screen again. Behind the scenes, the `map_to` method creates the missing page tables in the following way:
|
||
|
||
- Allocate an unused frame from the passed `frame_allocator`.
|
||
- Map the entry of the higher level table to that frame. Now the frame is accessible through the recursive page table.
|
||
- Zero the frame to create a new, empty page table.
|
||
- Continue with the next table level.
|
||
|
||
While our `create_mapping` function is just some example code, we are now able to create new mappings for arbitrary pages. This will be essential for allocating memory or implementing multithreading in future posts.
|
||
|
||
## Summary
|
||
|
||
In this post we learned how a recursive level 4 table entry can be used to map all page table frames to calculatable virtual addresses. We used this technique to implement an address translation function and to create a new mapping in the page tables.
|
||
|
||
We saw that the creation of new mappings requires unused frames for creating new page tables. Such a frame allocator can be implemented on top of the boot information structure that the bootloader passes to our kernel.
|
||
|
||
## What's next?
|
||
|
||
The next post will create a heap memory region for our kernel, which will allow us to [allocate memory] and use various [collection types].
|
||
|
||
[allocate memory]: https://doc.rust-lang.org/alloc/boxed/struct.Box.html
|
||
[collection types]: https://doc.rust-lang.org/alloc/collections/index.html
|