This is the fourth part of the Kernel booting process
where we will see first steps in protected mode, like checking that CPU supports long mode and SSE, paging, initializes the page tables and at the end we will discuss the transition to long mode.
NOTE: there will be much assembly code in this part, so if you are not familiar with that, you might want to consult a book about it
In the previous part we stopped at the jump to the 32-bit
entry point in arch/x86/boot/pmjump.S:
jmpl *%eax
You will recall that eax
register contains the address of the 32-bit entry point. We can read about this in the linux kernel x86 boot protocol:
When using bzImage, the protected-mode kernel was relocated to 0x100000
Let's make sure that it is true by looking at the register values at the 32-bit entry point:
eax 0x100000 1048576
ecx 0x0 0
edx 0x0 0
ebx 0x0 0
esp 0x1ff5c 0x1ff5c
ebp 0x0 0x0
esi 0x14470 83056
edi 0x0 0
eip 0x100000 0x100000
eflags 0x46 [ PF ZF ]
cs 0x10 16
ss 0x18 24
ds 0x18 24
es 0x18 24
fs 0x18 24
gs 0x18 24
We can see here that cs
register contains - 0x10
(as you may remember from the previous part, this is the second index in the Global Descriptor Table
), eip
register contains 0x100000
and the base address of all segments including the code segment are zero.
So we can get the physical address, it will be 0:0x100000
or just 0x100000
, as specified by the boot protocol. Now let's start with the 32-bit
entry point.
We can find the definition of the 32-bit
entry point in the arch/x86/boot/compressed/head_64.S assembly source code file:
__HEAD
.code32
ENTRY(startup_32)
....
....
....
ENDPROC(startup_32)
First of all, why the directory is named compressed
? Actually bzimage
is a gzipped vmlinux + header + kernel setup code
. We saw the kernel setup code in all of the previous parts. So, the main goal of the head_64.S
is to prepare for entering long mode, enter into it and then decompress the kernel. We will see all of the steps up to kernel decompression in this part.
You may find two files in the arch/x86/boot/compressed
directory:
but we will consider only head_64.S
source code file because, as you may remember, this book is only x86_64
related; Let's look at arch/x86/boot/compressed/Makefile. We can find the following make
target here:
vmlinux-objs-y := $(obj)/vmlinux.lds $(obj)/head_$(BITS).o $(obj)/misc.o \
$(obj)/string.o $(obj)/cmdline.o \
$(obj)/piggy.o $(obj)/cpuflags.o
Take a look on the $(obj)/head_$(BITS).o
.
This means that we will select which file to link based on what $(BITS)
is set to, either head_32.o
or head_64.o
. The $(BITS)
variable is defined elsewhere in arch/x86/Makefile based on the kernel configuration:
ifeq ($(CONFIG_X86_32),y)
BITS := 32
...
...
else
BITS := 64
...
...
endif
Now we know where to start, so let's do it.
As indicated above, we start in the arch/x86/boot/compressed/head_64.S assembly source code file. First we see the definition of the special section attribute before the startup_32
definition:
__HEAD
.code32
ENTRY(startup_32)
The __HEAD
is macro which is defined in include/linux/init.h header file and expands to the definition of the following section:
#define __HEAD .section ".head.text","ax"
with .head.text
name and ax
flags. In our case, these flags show us that this section is executable or in other words contains code. We can find definition of this section in the arch/x86/boot/compressed/vmlinux.lds.S linker script:
SECTIONS
{
. = 0;
.head.text : {
_head = . ;
HEAD_TEXT
_ehead = . ;
}
If you are not familiar with the syntax of GNU LD
linker scripting language, you can find more information in the documentation. In short, the .
symbol is a special variable of linker - location counter. The value assigned to it is an offset relative to the offset of the segment. In our case, we assign zero to location counter. This means that our code is linked to run from the 0
offset in memory. Moreover, we can find this information in comments:
Be careful parts of head_64.S assume startup_32 is at address 0.
Ok, now we know where we are, and now is the best time to look inside the startup_32
function.
In the beginning of the startup_32
function, we can see the cld
instruction which clears the DF
bit in the flags register. When direction flag is clear, all string operations like stos, scas and others will increment the index registers esi
or edi
. We need to clear direction flag because later we will use strings operations for clearing space for page tables, etc.
After we have cleared the DF
bit, next step is the check of the KEEP_SEGMENTS
flag from loadflags
kernel setup header field. If you remember we already saw loadflags
in the very first part of this book. There we checked CAN_USE_HEAP
flag to get ability to use heap. Now we need to check the KEEP_SEGMENTS
flag. This flag is described in the linux boot protocol documentation:
Bit 6 (write): KEEP_SEGMENTS
Protocol: 2.07+
- If 0, reload the segment registers in the 32bit entry point.
- If 1, do not reload the segment registers in the 32bit entry point.
Assume that %cs %ds %ss %es are all set to flat segments with
a base of 0 (or the equivalent for their environment).
So, if the KEEP_SEGMENTS
bit is not set in the loadflags
, we need to reset ds
, ss
and es
segment registers to a flat segment with base 0
. That we do:
testb $(1 << 6), BP_loadflags(%esi)
jnz 1f
cli
movl $(__BOOT_DS), %eax
movl %eax, %ds
movl %eax, %es
movl %eax, %ss
Remember that the __BOOT_DS
is 0x18
(index of data segment in the Global Descriptor Table). If KEEP_SEGMENTS
is set, we jump to the nearest 1f
label or update segment registers with __BOOT_DS
if it is not set. It is pretty easy, but here is one interesting moment. If you've read the previous part, you may remember that we already updated these segment registers right after we switched to protected mode in arch/x86/boot/pmjump.S. So why do we need to care about values of segment registers again? The answer is easy. The Linux kernel also has a 32-bit boot protocol and if a bootloader uses it to load the Linux kernel all code before the startup_32
will be missed. In this case, the startup_32
will be the first entry point of the Linux kernel right after the bootloader and there are no guarantees that segment registers will be in known state.
After we have checked the KEEP_SEGMENTS
flag and put the correct value to the segment registers, the next step is to calculate the difference between where we loaded and compiled to run. Remember that setup.ld.S
contains following definition: . = 0
at the start of the .head.text
section. This means that the code in this section is compiled to run from 0
address. We can see this in objdump
output:
arch/x86/boot/compressed/vmlinux: file format elf64-x86-64
Disassembly of section .head.text:
0000000000000000 <startup_32>:
0: fc cld
1: f6 86 11 02 00 00 40 testb $0x40,0x211(%rsi)
The objdump
util tells us that the address of the startup_32
is 0
but actually it's not so. Our current goal is to know where actually we are. It is pretty simple to do in long mode because it support rip
relative addressing but currently we are in protected mode. We will use common pattern to know the address of the startup_32
. We need to define a label and make a call to this label and pop the top of the stack to a register:
call label
label: pop %reg
After this, a register will contain the address of a label. Let's look at the similar code which searches address of the startup_32
in the Linux kernel:
leal (BP_scratch+4)(%esi), %esp
call 1f
1: popl %ebp
subl $1b, %ebp
As you remember from the previous part, the esi
register contains the address of the boot_params structure which was filled before we moved to the protected mode. The boot_params
structure contains a special field scratch
with offset 0x1e4
. These four bytes field will be temporary stack for call
instruction. We are getting the address of the scratch
field + 4 bytes and putting it in the esp
register. We add 4
bytes to the base of the BP_scratch
field because, as just described, it will be a temporary stack and the stack grows from top to down in x86_64
architecture. So our stack pointer will point to the top of the stack. Next, we can see the pattern that I've described above. We make a call to the 1f
label and put the address of this label to the ebp
register because we have return address on the top of stack after the call
instruction will be executed. So, for now we have an address of the 1f
label and now it is easy to get address of the startup_32
. We just need to subtract address of label from the address which we got from the stack:
startup_32 (0x0) +-----------------------+
| |
| |
| |
| |
| |
| |
| |
| |
1f (0x0 + 1f offset) +-----------------------+ %ebp - real physical address
| |
| |
+-----------------------+
startup_32
is linked to run at address 0x0
and this means that 1f
has the address 0x0 + offset to 1f
, approximately 0x21
bytes. The ebp
register contains the real physical address of the 1f
label. So, if we subtract 1f
from the ebp
we will get the real physical address of the startup_32
. The Linux kernel boot protocol describes that the base of the protected mode kernel is 0x100000
. We can verify this with gdb. Let's start the debugger and put breakpoint to the 1f
address, which is 0x100021
. If this is correct we will see 0x100021
in the ebp
register:
$ gdb
(gdb)$ target remote :1234
Remote debugging using :1234
0x0000fff0 in ?? ()
(gdb)$ br *0x100022
Breakpoint 1 at 0x100022
(gdb)$ c
Continuing.
Breakpoint 1, 0x00100022 in ?? ()
(gdb)$ i r
eax 0x18 0x18
ecx 0x0 0x0
edx 0x0 0x0
ebx 0x0 0x0
esp 0x144a8 0x144a8
ebp 0x100021 0x100021
esi 0x142c0 0x142c0
edi 0x0 0x0
eip 0x100022 0x100022
eflags 0x46 [ PF ZF ]
cs 0x10 0x10
ss 0x18 0x18
ds 0x18 0x18
es 0x18 0x18
fs 0x18 0x18
gs 0x18 0x18
If we execute the next instruction, subl $1b, %ebp
, we will see:
nexti
...
ebp 0x100000 0x100000
...
Ok, that's true. The address of the startup_32
is 0x100000
. After we know the address of the startup_32
label, we can prepare for the transition to long mode. Our next goal is to setup the stack and verify that the CPU supports long mode and SSE.
We could not setup the stack while we did not know the address of the startup_32
label. We can imagine the stack as an array and the stack pointer register esp
must point to the end of this array. Of course, we can define an array in our code, but we need to know its actual address to configure the stack pointer in a correct way. Let's look at the code:
movl $boot_stack_end, %eax
addl %ebp, %eax
movl %eax, %esp
The boot_stack_end
label, defined in the same arch/x86/boot/compressed/head_64.S assembly source code file and located in the .bss section:
.bss
.balign 4
boot_heap:
.fill BOOT_HEAP_SIZE, 1, 0
boot_stack:
.fill BOOT_STACK_SIZE, 1, 0
boot_stack_end:
First of all, we put the address of boot_stack_end
into the eax
register, so the eax
register contains the address of boot_stack_end
where it was linked, which is 0x0 + boot_stack_end
. To get the real address of boot_stack_end
, we need to add the real address of the startup_32
. As you remember, we have found this address above and put it to the ebp
register. In the end, the register eax
will contain real address of the boot_stack_end
and we just need to put to the stack pointer.
After we have set up the stack, next step is CPU verification. As we are going to execute transition to the long mode
, we need to check that the CPU supports long mode
and SSE
. We will do it by the call of the verify_cpu
function:
call verify_cpu
testl %eax, %eax
jnz no_longmode
This function defined in the arch/x86/kernel/verify_cpu.S assembly file and just contains a couple of calls to the cpuid instruction. This instruction is used for getting information about the processor. In our case, it checks long mode
and SSE
support and returns 0
on success or 1
on fail in the eax
register.
If the value of the eax
is not zero, we jump to the no_longmode
label which just stops the CPU by the call of the hlt
instruction while no hardware interrupt will not happen:
no_longmode:
1:
hlt
jmp 1b
If the value of the eax
register is zero, everything is ok and we are able to continue.
The next step is calculating relocation address for decompression if needed. First, we need to know what it means for a kernel to be relocatable
. We already know that the base address of the 32-bit entry point of the Linux kernel is 0x100000
, but that is a 32-bit entry point. The default base address of the Linux kernel is determined by the value of the CONFIG_PHYSICAL_START
kernel configuration option. Its default value is 0x1000000
or 16 MB
. The main problem here is that if the Linux kernel crashes, a kernel developer must have a rescue kernel
for kdump which is configured to load from a different address. The Linux kernel provides special configuration option to solve this problem: CONFIG_RELOCATABLE
. As we can read in the documentation of the Linux kernel:
This builds a kernel image that retains relocation information
so it can be loaded someplace besides the default 1MB.
Note: If CONFIG_RELOCATABLE=y, then the kernel runs from the address
it has been loaded at and the compile time physical address
(CONFIG_PHYSICAL_START) is used as the minimum location.
In simple terms, this means that the Linux kernel with the same configuration can be booted from different addresses. Technically, this is done by compiling the decompressor as position independent code. If we look at arch/x86/boot/compressed/Makefile, we will see that the decompressor is indeed compiled with the -fPIC
flag:
KBUILD_CFLAGS += -fno-strict-aliasing -fPIC
When we are using position-independent code an address is obtained by adding the address field of the command and the value of the program counter. We can load code which uses such addressing from any address. That's why we had to get the real physical address of startup_32
. Now let's get back to the Linux kernel code. Our current goal is to calculate an address where we can relocate the kernel for decompression. Calculation of this address depends on CONFIG_RELOCATABLE
kernel configuration option. Let's look at the code:
#ifdef CONFIG_RELOCATABLE
movl %ebp, %ebx
movl BP_kernel_alignment(%esi), %eax
decl %eax
addl %eax, %ebx
notl %eax
andl %eax, %ebx
cmpl $LOAD_PHYSICAL_ADDR, %ebx
jge 1f
#endif
movl $LOAD_PHYSICAL_ADDR, %ebx
1:
addl $z_extract_offset, %ebx
Remember that the value of the ebp
register is the physical address of the startup_32
label. If the CONFIG_RELOCATABLE
kernel configuration option is enabled during kernel configuration, we put this address in the ebx
register, align it to a multiple of 2MB
and compare it with the LOAD_PHYSICAL_ADDR
value. The LOAD_PHYSICAL_ADDR
macro is defined in the arch/x86/include/asm/boot.h header file and it looks like this:
#define LOAD_PHYSICAL_ADDR ((CONFIG_PHYSICAL_START \
+ (CONFIG_PHYSICAL_ALIGN - 1)) \
& ~(CONFIG_PHYSICAL_ALIGN - 1))
As we can see it just expands to the aligned CONFIG_PHYSICAL_ALIGN
value which represents the physical address of where to load the kernel. After comparison of the LOAD_PHYSICAL_ADDR
and value of the ebx
register, we add the offset from the startup_32
where to decompress the compressed kernel image. If the CONFIG_RELOCATABLE
option is not enabled during kernel configuration, we just put the default address where to load kernel and add z_extract_offset
to it.
After all of these calculations, we will have ebp
which contains the address where we loaded it and ebx
set to the address of where kernel will be moved after decompression.
When we have the base address where we will relocate the compressed kernel image, we need to do one last step before we can transition to 64-bit mode. First, we need to update the Global Descriptor Table:
leal gdt(%ebp), %eax
movl %eax, gdt+2(%ebp)
lgdt gdt(%ebp)
Here we put the base address from ebp
register with gdt
offset into the eax
register. Next we put this address into ebp
register with offset gdt+2
and load the Global Descriptor Table
with the lgdt
instruction. To understand the magic with gdt
offsets we need to look at the definition of the Global Descriptor Table
. We can find its definition in the same source code file:
.data
gdt:
.word gdt_end - gdt
.long gdt
.word 0
.quad 0x0000000000000000 /* NULL descriptor */
.quad 0x00af9a000000ffff /* __KERNEL_CS */
.quad 0x00cf92000000ffff /* __KERNEL_DS */
.quad 0x0080890000000000 /* TS descriptor */
.quad 0x0000000000000000 /* TS continued */
gdt_end:
We can see that it is located in the .data
section and contains five descriptors: null
descriptor, kernel code segment, kernel data segment and two task descriptors. We already loaded the Global Descriptor Table
in the previous part, and now we're doing almost the same here, but descriptors with CS.L = 1
and CS.D = 0
for execution in 64
bit mode. As we can see, the definition of the gdt
starts from two bytes: gdt_end - gdt
which represents the last byte in the gdt
table or table limit. The next four bytes contains base address of the gdt
. Remember that the Global Descriptor Table
is stored in the 48-bits GDTR
which consists of two parts:
- size(16-bit) of global descriptor table;
- address(32-bit) of the global descriptor table.
So, we put address of the gdt
to the eax
register and then we put it to the .long gdt
or gdt+2
in our assembly code. From now we have formed structure for the GDTR
register and can load the Global Descriptor Table
with the lgtd
instruction.
After we have loaded the Global Descriptor Table
, we must enable PAE mode by putting the value of the cr4
register into eax
, setting 5 bit in it and loading it again into cr4
:
movl %cr4, %eax
orl $X86_CR4_PAE, %eax
movl %eax, %cr4
Now we are almost finished with all preparations before we can move into 64-bit mode. The last step is to build page tables, but before that, here is some information about long mode.
Long mode is the native mode for x86_64 processors. First, let's look at some differences between x86_64
and the x86
.
The 64-bit
mode provides features such as:
- New 8 general purpose registers from
r8
tor15
+ all general purpose registers are 64-bit now; - 64-bit instruction pointer -
RIP
; - New operating mode - Long mode;
- 64-Bit Addresses and Operands;
- RIP Relative Addressing (we will see an example of it in the next parts).
Long mode is an extension of legacy protected mode. It consists of two sub-modes:
- 64-bit mode;
- compatibility mode.
To switch into 64-bit
mode we need to do following things:
- Enable PAE;
- Build page tables and load the address of the top level page table into the
cr3
register; - Enable
EFER.LME
; - Enable paging.
We already enabled PAE
by setting the PAE
bit in the cr4
control register. Our next goal is to build the structure for paging. We will see this in next paragraph.
So, we already know that before we can move into 64-bit
mode, we need to build page tables, so, let's look at the building of early 4G
boot page tables.
NOTE: I will not describe the theory of virtual memory here. If you need to know more about it, see links at the end of this part.
The Linux kernel uses 4-level
paging, and we generally build 6 page tables:
- One
PML4
orPage Map Level 4
table with one entry; - One
PDP
orPage Directory Pointer
table with four entries; - Four Page Directory tables with a total of
2048
entries.
Let's look at the implementation of this. First of all, we clear the buffer for the page tables in memory. Every table is 4096
bytes, so we need clear 24
kilobyte buffer:
leal pgtable(%ebx), %edi
xorl %eax, %eax
movl $((4096*6)/4), %ecx
rep stosl
We put the address of pgtable
plus ebx
(remember that ebx
contains the address to relocate the kernel for decompression) in the edi
register, clear the eax
register and set the ecx
register to 6144
. The rep stosl
instruction will write the value of the eax
to edi
, increase value of the edi
register by 4
and decrease the value of the ecx
register by 1
. This operation will be repeated while the value of the ecx
register is greater than zero. That's why we put 6144
in ecx
.
pgtable
is defined at the end of arch/x86/boot/compressed/head_64.S assembly file and is:
.section ".pgtable","a",@nobits
.balign 4096
pgtable:
.fill 6*4096, 1, 0
As we can see, it is located in the .pgtable
section and its size is 24
kilobytes.
After we have got buffer for the pgtable
structure, we can start to build the top level page table - PML4
- with:
leal pgtable + 0(%ebx), %edi
leal 0x1007 (%edi), %eax
movl %eax, 0(%edi)
Here again, we put the address of the pgtable
relative to ebx
or in other words relative to address of the startup_32
to the edi
register. Next, we put this address with offset 0x1007
in the eax
register. The 0x1007
is 4096
bytes which is the size of the PML4
plus 7
. The 7
here represents flags of the PML4
entry. In our case, these flags are PRESENT+RW+USER
. In the end, we just write first the address of the first PDP
entry to the PML4
.
In the next step we will build four Page Directory
entries in the Page Directory Pointer
table with the same PRESENT+RW+USE
flags:
leal pgtable + 0x1000(%ebx), %edi
leal 0x1007(%edi), %eax
movl $4, %ecx
1: movl %eax, 0x00(%edi)
addl $0x00001000, %eax
addl $8, %edi
decl %ecx
jnz 1b
We put the base address of the page directory pointer which is 4096
or 0x1000
offset from the pgtable
table in edi
and the address of the first page directory pointer entry in eax
register. Put 4
in the ecx
register, it will be a counter in the following loop and write the address of the first page directory pointer table entry to the edi
register. After this edi
will contain the address of the first page directory pointer entry with flags 0x7
. Next we just calculate the address of following page directory pointer entries where each entry is 8
bytes, and write their addresses to eax
. The last step of building paging structure is the building of the 2048
page table entries with 2-MByte
pages:
leal pgtable + 0x2000(%ebx), %edi
movl $0x00000183, %eax
movl $2048, %ecx
1: movl %eax, 0(%edi)
addl $0x00200000, %eax
addl $8, %edi
decl %ecx
jnz 1b
Here we do almost the same as in the previous example, all entries will be with flags - $0x00000183
- PRESENT + WRITE + MBZ
. In the end, we will have 2048
pages with 2-MByte
page or:
>>> 2048 * 0x00200000
4294967296
4G
page table. We just finished to build our early page table structure which maps 4
gigabytes of memory and now we can put the address of the high-level page table - PML4
- in cr3
control register:
leal pgtable(%ebx), %eax
movl %eax, %cr3
That's all. All preparation are finished and now we can see transition to the long mode.
First of all we need to set the EFER.LME
flag in the MSR to 0xC0000080
:
movl $MSR_EFER, %ecx
rdmsr
btsl $_EFER_LME, %eax
wrmsr
Here we put the MSR_EFER
flag (which is defined in arch/x86/include/uapi/asm/msr-index.h) in the ecx
register and call rdmsr
instruction which reads the MSR register. After rdmsr
executes, we will have the resulting data in edx:eax
which depends on the ecx
value. We check the EFER_LME
bit with the btsl
instruction and write data from eax
to the MSR
register with the wrmsr
instruction.
In the next step, we push the address of the kernel segment code to the stack (we defined it in the GDT) and put the address of the startup_64
routine in eax
.
pushl $__KERNEL_CS
leal startup_64(%ebp), %eax
After this we push this address to the stack and enable paging by setting PG
and PE
bits in the cr0
register:
movl $(X86_CR0_PG | X86_CR0_PE), %eax
movl %eax, %cr0
and execute:
lret
instruction. Remember that we pushed the address of the startup_64
function to the stack in the previous step, and after the lret
instruction, the CPU extracts the address of it and jumps there.
After all of these steps we're finally in 64-bit mode:
.code64
.org 0x200
ENTRY(startup_64)
....
....
....
That's all!
This is the end of the fourth part linux kernel booting process. If you have questions or suggestions, ping me in twitter 0xAX, drop me email or just create an issue.
In the next part, we will see kernel decompression and much more.
Please note that English is not my first language and I am really sorry for any inconvenience. If you find any mistakes please send me PR to linux-insides.