I tried to explain everything in detail and to keep the code as simple as possible. If you have any questions, suggestions, or issues, please leave a comment or on Github. The source code is available in a repository, too.

To avoid bugs and strange errors on old CPUs we should check if the processor supports every needed feature. If not, the kernel should abort and display an error message. To handle errors easily, we create an error procedure in boot.asm. It prints a rudimentary ERR: X message, where X is an error code letter, and hangs:

At address 0xb8000 begins the so-called VGA text buffer. It’s an array of screen characters that are displayed by the graphics card. A will cover the VGA buffer in detail and create a Rust interface to it. But for now, manual bit-fiddling is the easiest option.

A screen character consists of a 8 bit color code and a 8 bit ASCII character. We used the color code 4f for all characters, which means white text on red background. 0x52 is an ASCII R, 0x45 is an E, 0x3a is a :, and 0x20 is a space. The second space is overwritten by the given ASCII byte. Finally the CPU is stopped with the hlt instruction.

Now we can add some check functions. A function is just a normal label with an ret (return) instruction at the end. The call instruction can be used to call it. Unlike the jmp instruction that just jumps to a memory address, the call instruction will push a return address to the stack (and the ret will jump to this address). But we don’t have a stack yet. The in the esp register could point to some important data or even invalid memory. So we need to update it and point it to some valid stack memory.

To create stack memory we reserve some bytes at the end of our boot.asm:

...
section .bss
stack_bottom:
    resb 64
stack_top:

A stack doesn’t need to be initialized because we will pop only when we pushed before. So storing the stack memory in the executable file would make it unnecessary large. By using the section and the resb (reserve byte) command, we just store the length of the uninitialized data (= 64). When loading the executable, GRUB will create the section of required size in memory.

To use the new stack, we update the stack pointer register right after start:

global start
section .text
bits 32
start:
    mov esp, stack_top
    ; print `OK` to screen
    ...

We use stack_top because the stack grows downwards: A push eax subtracts 4 from esp and does a mov [esp], eax afterwards (eax is a general purpose register).

Now we have a valid stack pointer and are able to call functions. The following check functions are just here for completeness and I won’t explain details. Basically they all work the same: They will check for a feature and jump to error if it’s not available.

Multiboot check

We rely on some Multiboot features in the next posts. To make sure the kernel was really loaded by a Multiboot compliant bootloader, we can check the eax register. According to the Multiboot specification (), the bootloader must write the magic value 0x36d76289 to it before loading a kernel. To verify that we can add a simple function:

check_multiboot:
    cmp eax, 0x36d76289
    jne .no_multiboot
    ret
.no_multiboot:
    mov al, "0"
    jmp error

We use the cmp instruction to compare the value in eax to the magic value. If the values are equal, the cmp instruction sets the zero flag in the FLAGS register. The jne (“jump if not equal”) instruction reads this zero flag and jumps to the given address if it’s not set. Thus we jump to the .no_multiboot label if eax does not contain the magic value.

In no_multiboot, we use the jmp (“jump”) instruction to jump to our error function. We could just as well use the call instruction, which additionally pushes the return address. But the return address is not needed because error never returns. To pass 0 as error code to the error function, we move it into al before the jump (error will read it from there).

CPUID is a CPU instruction that can be used to get various information about the CPU. But not every processor supports it. CPUID detection is quite laborious, so we just copy a detection function from the :

check_cpuid:
    ; Check if CPUID is supported by attempting to flip the ID bit (bit 21)
    ; in the FLAGS register. If we can flip it, CPUID is available.
    ; Copy FLAGS in to EAX via stack
    pushfd
    pop eax
    ; Copy to ECX as well for comparing later on
    mov ecx, eax
    ; Flip the ID bit
    xor eax, 1 << 21
    ; Copy EAX to FLAGS via the stack
    push eax
    popfd
    ; Copy FLAGS back to EAX (with the flipped bit if CPUID is supported)
    pushfd
    pop eax
    ; Restore FLAGS from the old version stored in ECX (i.e. flipping the
    ; ID bit back if it was ever flipped).
    push ecx
    popfd
    ; Compare EAX and ECX. If they are equal then that means the bit
    ; wasn't flipped, and CPUID isn't supported.
    cmp eax, ecx
    je .no_cpuid
    ret
.no_cpuid:
    mov al, "1"
    jmp error

Basically, the CPUID instruction is supported if we can flip some bit in the FLAGS register. We can’t operate on the flags register directly, so we need to load it into some general purpose register such as eax first. The only way to do this is to push the FLAGS register on the stack through the pushfd instruction and then pop it into eax. Equally, we write it back through push ecx and popfd. To flip the bit we use the xor instruction to perform an . Finally we compare the two values and jump to .no_cpuid if both are equal (je – “jump if equal”). The .no_cpuid code just jumps to the error function with error code 1.

Don’t worry, you don’t need to understand the details.

Long Mode check

Now we can use CPUID to detect whether long mode can be used. I use code from again:

    ; test if extended processor info in available
    mov eax, 0x80000000    ; implicit argument for cpuid
    cpuid                  ; get highest supported argument
    jb .no_long_mode       ; if it's less, the CPU is too old for long mode
    ; use extended info to test if long mode is available
    mov eax, 0x80000001    ; argument for extended processor info
    cpuid                  ; returns various feature bits in ecx and edx
    test edx, 1 << 29      ; test if the LM-bit is set in the D-register
    jz .no_long_mode       ; If it's not set, there is no long mode
    ret
.no_long_mode:
    mov al, "2"
    jmp error

Like many low-level things, CPUID is a bit strange. Instead of taking a parameter, the cpuid instruction implicitly uses the eax register as argument. To test if long mode is available, we need to call cpuid with 0x80000001 in eax. This loads some information to the ecx and edx registers. Long mode is supported if the 29th bit in edx is set. Wikipedia has detailed information.

If you look at the assembly above, you’ll probably notice that we call cpuid twice. The reason is that the CPUID command started with only a few functions and was extended over time. So old processors may not know the 0x80000001 argument at all. To test if they do, we need to invoke cpuid with 0x80000000 in eax first. It returns the highest supported parameter value in eax. If it’s at least 0x80000001, we can test for long mode as described above. Else the CPU is old and doesn’t know what long mode is either. In that case, we directly jump to .no_long_mode through the jb instruction (“jump if below”).

We just call these check functions right after start:

When the CPU doesn’t support a needed feature, we get an error message with an unique error code. Now we can start the real work.

In long mode, x86 uses a page size of 4096 bytes and a 4 level page table that consists of:

• the Page-Map Level-4 Table (PML4),
• the Page-Directory Pointer Table (PDP),
• the Page-Directory Table (PD),
• and the Page Table (PT).

As I don’t like these names, I will call them P4, P3, P2, and P1 from now on.

Each page table contains 512 entries and one entry is 8 bytes, so they fit exactly in one page (512*8 = 4096). To translate a virtual address to a physical address the CPU[^hardware_lookup] will do the following[^virtual_physical_translation_source]:

1. Get the address of the P4 table from the CR3 register
2. Use bits 39-47 (9 bits) as an index into P4 (2^9 = 512 = number of entries)
3. Use the following 9 bits as an index into P3
4. Use the following 9 bits as an index into P2
5. Use the following 9 bits as an index into P1
6. Use the last 12 bits as page offset (2^12 = 4096 = page size)

But what happens to bits 48-63 of the 64-bit virtual address? Well, they can’t be used. The “64-bit” long mode is in fact just a 48-bit mode. The bits 48-63 must be copies of bit 47, so each valid virtual address is still unique. For more information see .

An entry in the P4, P3, P2, and P1 tables consists of the page aligned 52-bit physical address of the frame or the next page table and the following bits that can be OR-ed in:

Set Up Identity Paging

When we switch to long mode, paging will be activated automatically. The CPU will then try to read the instruction at the following address, but this address is now a virtual address. So we need to do identity mapping, i.e. map a physical address to the same virtual address.

The huge page bit is now very useful to us. It creates a 2MiB (when used in P2) or even a 1GiB page (when used in P3). So we could map the first gigabytes of the kernel with only one P4 and one P3 table by using 1GiB pages. Unfortunately 1GiB pages are relatively new feature, for example Intel introduced it 2010 in the . Therefore we will use 2MiB pages instead to make our kernel compatible to older computers, too.

To identity map the first gigabyte of our kernel with 512 2MiB pages, we need one P4, one P3, and one P2 table. Of course we will replace them with finer-grained tables later. But now that we’re stuck with assembly, we choose the easiest way.

We can add these two tables at the beginning[^page_table_alignment] of the .bss section:

...
section .bss
align 4096
p4_table:
    resb 4096
p3_table:
    resb 4096
p2_table:
    resb 4096
stack_bottom:
    resb 64
stack_top:

The resb command reserves the specified amount of bytes without initializing them, so the 8KiB don’t need to be saved in the executable. The align 4096 ensures that the page tables are page aligned.

When GRUB creates the .bss section in memory, it will initialize it to 0. So the p4_table is already valid (it contains 512 non-present entries) but not very useful. To be able to map 2MiB pages, we need to link P4’s first entry to the p3_table and P3’s first entry to the the p2_table:

set_up_page_tables:
    ; map first P4 entry to P3 table
    mov eax, p3_table
    or eax, 0b11 ; present + writable
    mov [p4_table], eax
    ; map first P3 entry to P2 table
    mov eax, p2_table
    or eax, 0b11 ; present + writable
    mov [p3_table], eax
    ; TODO map each P2 entry to a huge 2MiB page
    ret

We just set the present and writable bits (0b11 is a binary number) in the aligned P3 table address and move it to the first 4 bytes of the P4 table. Then we do the same to link the first P3 entry to the p2_table.

Now we need to map P2’s first entry to a huge page starting at 0, P2’s second entry to a huge page starting at 2MiB, P2’s third entry to a huge page starting at 4MiB, and so on. It’s time for our first (and only) assembly loop:

set_up_page_tables:
    ...
    ; map each P2 entry to a huge 2MiB page
    mov ecx, 0         ; counter variable
.map_p2_table:
    ; map ecx-th P2 entry to a huge page that starts at address 2MiB*ecx
    mov eax, 0x200000  ; 2MiB
    mul ecx            ; start address of ecx-th page
    or eax, 0b10000011 ; present + writable + huge
    mov [p2_table + ecx * 8], eax ; map ecx-th entry
    inc ecx            ; increase counter
    cmp ecx, 512       ; if counter == 512, the whole P2 table is mapped
    jne .map_p2_table  ; else map the next entry
    ret

Maybe I should first explain how an assembly loop works. We use the ecx register as a counter variable, just like i in a for loop. After mapping the ecx-th entry, we increase ecx by one and jump to .map_p2_table again if it’s still smaller than 512.

To map a P2 entry we first calculate the start address of its page in eax: The ecx-th entry needs to be mapped to ecx * 2MiB. We use the mul operation for that, which multiplies eax with the given register and stores the result in eax. Then we set the present, writable, and huge page bits and write it to the P2 entry. The address of the ecx-th entry in P2 is p2_table + ecx * 8, because each entry is 8 bytes large.

Now the first gigabyte (512 * 2MiB) of our kernel is identity mapped and thus accessible through the same physical and virtual addresses.

To enable paging and enter long mode, we need to do the following:

1. write the address of the P4 table to the CR3 register (the CPU will look there, see the )
2. long mode is an extension of Physical Address Extension (PAE), so we need to enable PAE first
3. Set the long mode bit in the EFER register
4. Enable Paging

The assembly function looks like this (some boring bit-moving to various registers):

enable_paging:
    ; load P4 to cr3 register (cpu uses this to access the P4 table)
    mov eax, p4_table
    mov cr3, eax
    mov eax, cr4
    or eax, 1 << 5
    ; set the long mode bit in the EFER MSR (model specific register)
    mov ecx, 0xC0000080
    rdmsr
    or eax, 1 << 8
    wrmsr
    ; enable paging in the cr0 register
    mov eax, cr0
    or eax, 1 << 31
    mov cr0, eax
    ret

The or eax, 1 << X is a common pattern. It sets the bit X in the eax register (<< is a left shift). Through rdmsr and wrmsr it’s possible to read/write to the so-called model specific registers at address ecx (in this case ecx points to the EFER register).

Finally we need to call our new functions in start:

...
start:
    mov esp, stack_top
    call check_multiboot
    call check_cpuid
    call check_long_mode
    call set_up_page_tables ; new
    call enable_paging     ; new
    ; print `OK` to screen
    mov dword [0xb8000], 0x2f4b2f4f
    hlt
...

After enabling Paging, the processor is in long mode. So we can use 64-bit instructions now, right? Wrong. The processor is still in a 32-bit compatibility submode. To actually execute 64-bit code, we need to set up a new Global Descriptor Table. The Global Descriptor Table (GDT) was used for Segmentation in old operating systems. I won’t explain Segmentation but the Three Easy Pieces OS book has good introduction () again.

Today almost everyone uses Paging instead of Segmentation (and so do we). But on x86, a GDT is always required, even when you’re not using Segmentation. GRUB has set up a valid 32-bit GDT for us but now we need to switch to a long mode GDT.

A GDT always starts with a 0-entry and contains an arbitrary number of segment entries afterwards. A 64-bit entry has the following format:

We need one code segment, a data segment is not necessary in 64-bit mode. Code segments have the following bits set: descriptor type, present, executable and the 64-bit flag. Translated to assembly the long mode GDT looks like this:

We chose the .rodata section here because it’s initialized read-only data. The dq command stands for define quad and outputs a 64-bit constant (similar to dw and dd). And the (1<<43) is a bit shift that sets bit 43.

Loading the GDT

To load our new 64-bit GDT, we have to tell the CPU its address and length. We do this by passing the memory location of a special pointer structure to the lgdt (load GDT) instruction. The pointer structure looks like this:

gdt64:
    dq 0 ; zero entry
    dq (1<<43) | (1<<44) | (1<<47) | (1<<53) ; code segment
.pointer:
    dw $ - gdt64 - 1
    dq gdt64

The first 2 bytes specify the (GDT length - 1). The $ is a special symbol that is replaced with the current address (it’s equal to .pointer in our case). The following 8 bytes specify the GDT address. Labels that start with a point (such as .pointer) are sub-labels of the last label without point. To access them, they must be prefixed with the parent label (e.g., gdt64.pointer).

Now we can load the GDT in start:

start:
    ...
    call enable_paging
    ; load the 64-bit GDT
    lgdt [gdt64.pointer]
    ; print `OK` to screen
    ...

When you still see the green OK, everything went fine and the new GDT is loaded. But we still can’t execute 64-bit code: The code selector register cs still has the values from the old GDT. To update it, we need to load it with the GDT offset (in bytes) of the desired segment. In our case the code segment starts at byte 8 of the GDT, but we don’t want to hardcode that 8 (in case we modify our GDT later). Instead, we add a .code label to our GDT, that calculates the offset directly from the GDT:

section .rodata
gdt64:
    dq 0 ; zero entry
.code: equ $ - gdt64 ; new
    dq (1<<43) | (1<<44) | (1<<47) | (1<<53) ; code segment
.pointer:
    ...

We can’t just use a normal label here, since we need the table offset. We calculate this offset using the current address $ and set the label to this value using equ. Now we can use gdt64.code instead of 8 and this label will still work if we modify the GDT.

In order to finally enter the true 64-bit mode, we need to load cs with gdt64.code. But we can’t do it through mov. The only way to reload the code selector is a far jump or a far return. These instructions work like a normal jump/return but change the code selector. We use a far jump to a long mode label:

global start
extern long_mode_start
...
start:
    ...
    lgdt [gdt64.pointer]
    jmp gdt64.code:long_mode_start
...

The actual long_mode_start label is defined as extern, so it’s part of another file. The jmp gdt64.code:long_mode_start is the mentioned far jump.

I put the 64-bit code into a new file to separate it from the 32-bit code, thereby we can’t call the (now invalid) 32-bit code accidentally. The new file (I named it long_mode_init.asm) looks like this:

global long_mode_start
section .text
bits 64
long_mode_start:
    ; print `OKAY` to screen
    mov rax, 0x2f592f412f4b2f4f
    mov qword [0xb8000], rax
    hlt

You should see a green OKAY on the screen. Some notes on this last step:

• As the CPU expects 64-bit instructions now, we use bits 64
• We can now use the extended registers. Instead of the 32-bit eax, ebx, etc. we now have the 64-bit rax, rbx, …

Congratulations! You have successfully wrestled through this CPU configuration and compatibility mode mess :).

One Last Thing

Above, we reloaded the code segment register cs with the new GDT offset. However, the data segment registers ss, ds, es, fs, and still contain the data segment offsets of the old GDT. This isn’t necessarily bad, since they’re ignored by almost all instructions in 64-bit mode. However, there are a few instructions that expect a valid data segment descriptor or the null descriptor in those registers. An example is the the iretq instruction that we’ll need in the post.

To avoid future problems, we reload all data segment registers with null:

It’s time to finally leave assembly behind and switch to . Rust is a systems language without garbage collections that guarantees memory safety. Through a real type system and many abstractions it feels like a high-level language but can still be low-level enough for OS development. The next post describes the Rust setup.

[^hardwarelookup]: In the x86 architecture, the page tables are _hardware walked, so the CPU will look at the table on its own when it needs a translation. Other architectures, for example MIPS, just throw an exception and let the OS translate the virtual address.

[^virtual_physical_translation_source]: Image source: Wikipedia, with modified font size, page table naming, and removed sign extended bits. The modified file is licensed under the Creative Commons Attribution-Share Alike 3.0 Unported license.

[^page_table_alignment]: Page tables need to be page-aligned as the bits 0-11 are used for flags. By putting these tables at the beginning of .bss, the linker can just page align the whole section and we don’t have unused padding bytes in between.