This chapter is an overview of the boot and system initialization processes, starting from the BIOS (firmware) POST, to the first user process creation. Since the initial steps of system startup are very architecture dependent, the IA-32 architecture is used as an example. But the AMD64 and ARM64 architectures are much more important and compelling examples and should be explained in the near future according to the topic of this document.

The FreeBSD boot process can be surprisingly complex. After control is passed from the BIOS, a considerable amount of low-level configuration must be done before the kernel can be loaded and executed. This setup must be done in a simple and flexible manner, allowing the user a great deal of customization possibilities.

The boot process is an extremely machine-dependent activity. Not only must code be written for every computer architecture, but there may also be multiple types of booting on the same architecture. For example, a directory listing of stand reveals a great amount of architecture-dependent code. There is a directory for each of the various supported architectures. FreeBSD supports the CSM boot standard (Compatibility Support Module). So CSM is supported (with both GPT and MBR partitioning support) and UEFI booting (GPT is totally supported, MBR is mostly supported). It also supports loading files from ext2fs, MSDOS, UFS and ZFS. FreeBSD also supports the boot environment feature of ZFS which allows the HOST OS to communicate details about what to boot that go beyond a simple partition as was possible in the past. But UEFI is more relevant than the CSM these days. The example that follows shows booting an x86 computer from an MBR-partitioned hard drive with the FreeBSD boot0 multi-boot loader stored in the very first sector. That boot code starts the FreeBSD three-stage boot process.

The key to understanding this process is that it is a series of stages of increasing complexity. These stages are boot1, boot2, and loader (see boot(8) for more detail). The boot system executes each stage in sequence. The last stage, loader, is responsible for loading the FreeBSD kernel. Each stage is examined in the following sections.

Here is an example of the output generated by the different boot stages. Actual output may differ from machine to machine:

When the computer powers on, the processor’s registers are set to some predefined values. One of the registers is the instruction pointer register, and its value after a power on is well defined: it is a 32-bit value of 0xfffffff0. The instruction pointer register (also known as the Program Counter) points to code to be executed by the processor. Another important register is the cr0 32-bit control register, and its value just after a reboot is 0. One of cr0's bits, the PE (Protection Enabled) bit, indicates whether the processor is running in 32-bit protected mode or 16-bit real mode. Since this bit is cleared at boot time, the processor boots in 16-bit real mode. Real mode means, among other things, that linear and physical addresses are identical. The reason for the processor not to start immediately in 32-bit protected mode is backwards compatibility. In particular, the boot process relies on the services provided by the BIOS, and the BIOS itself works in legacy, 16-bit code.

The value of 0xfffffff0 is slightly less than 4 GB, so unless the machine has 4 GB of physical memory, it cannot point to a valid memory address. The computer’s hardware translates this address so that it points to a BIOS memory block.

The BIOS (Basic Input Output System) is a chip on the motherboard that has a relatively small amount of read-only memory (ROM). This memory contains various low-level routines that are specific to the hardware supplied with the motherboard. The processor will first jump to the address 0xfffffff0, which really resides in the BIOS’s memory. Usually this address contains a jump instruction to the BIOS’s POST routines.

The POST (Power On Self Test) is a set of routines including the memory check, system bus check, and other low-level initialization so the CPU can set up the computer properly. The important step of this stage is determining the boot device. Modern BIOS implementations permit the selection of a boot device, allowing booting from a floppy, CD-ROM, hard disk, or other devices.

The very last thing in the POST is the INT 0x19 instruction. The INT 0x19 handler reads 512 bytes from the first sector of boot device into the memory at address 0x7c00. The term first sector originates from hard drive architecture, where the magnetic plate is divided into a number of cylindrical tracks. Tracks are numbered, and every track is divided into a number (usually 64) of sectors. Track numbers start at 0, but sector numbers start from 1. Track 0 is the outermost on the magnetic plate, and sector 1, the first sector, has a special purpose. It is also called the MBR, or Master Boot Record. The remaining sectors on the first track are never used.

This sector is our boot-sequence starting point. As we will see, this sector contains a copy of our boot0 program. A jump is made by the BIOS to address 0x7c00 so it starts executing.

After control is received from the BIOS at memory address 0x7c00, boot0 starts executing. It is the first piece of code under FreeBSD control. The task of boot0 is quite simple: scan the partition table and let the user choose which partition to boot from. The Partition Table is a special, standard data structure embedded in the MBR (hence embedded in boot0) describing the four standard PC "partitions". boot0 resides in the filesystem as /boot/boot0. It is a small 512-byte file, and it is exactly what FreeBSD’s installation procedure wrote to the hard disk’s MBR if you chose the "bootmanager" option at installation time. Indeed, boot0 is the MBR.

As mentioned previously, we’re calling the BIOS INT 0x19 to load the MBR (boot0) into memory at address 0x7c00. The source file for boot0 can be found in stand/i386/boot0/boot0.S - which is an awesome piece of code written by Robert Nordier.

A special structure starting from offset 0x1be in the MBR is called the partition table. It has four records of 16 bytes each, called partition records, which represent how the hard disk is partitioned, or, in FreeBSD’s terminology, sliced. One byte of those 16 says whether a partition (slice) is bootable or not. Exactly one record must have that flag set, otherwise boot0's code will refuse to proceed.

A partition record has the following fields:

A partition record descriptor contains information about where exactly the partition resides on the drive. Both descriptors, LBA and CHS, describe the same information, but in different ways: LBA (Logical Block Addressing) has the starting sector for the partition and the partition’s length, while CHS (Cylinder Head Sector) has coordinates for the first and last sectors of the partition. The partition table ends with the special signature 0xaa55.

The MBR must fit into 512 bytes, a single disk sector. This program uses low-level "tricks" like taking advantage of the side effects of certain instructions and reusing register values from previous operations to make the most out of the fewest possible instructions. Care must also be taken when handling the partition table, which is embedded in the MBR itself. For these reasons, be very careful when modifying boot0.S.

Note that the boot0.S source file is assembled "as is": instructions are translated one by one to binary, with no additional information (no ELF file format, for example). This kind of low-level control is achieved at link time through special control flags passed to the linker. For example, the text section of the program is set to be located at address 0x600. In practice this means that boot0 must be loaded to memory address 0x600 in order to function properly.

It is worth looking at the Makefile for boot0 (stand/i386/boot0/Makefile), as it defines some of the run-time behavior of boot0. For instance, if a terminal connected to the serial port (COM1) is used for I/O, the macro SIO must be defined (-DSIO). -DPXE enables boot through PXE by pressing F6. Additionally, the program defines a set of flags that allow further modification of its behavior. All of this is illustrated in the Makefile. For example, look at the linker directives which command the linker to start the text section at address 0x600, and to build the output file "as is" (strip out any file formatting):

Let us now start our study of the MBR, or boot0, starting where execution begins.

This first block of code is the entry point of the program. It is where the BIOS transfers control. First, it makes sure that the string operations autoincrement its pointer operands (the cld instruction) ^[2]. Then, as it makes no assumption about the state of the segment registers, it initializes them. Finally, it sets the stack pointer register (%sp) to ($LOAD = address 0x7c00), so we have a working stack.

The next block is responsible for the relocation and subsequent jump to the relocated code.

As boot0 is loaded by the BIOS to address 0x7C00, it copies itself to address 0x600 and then transfers control there (recall that it was linked to execute at address 0x600). The source address, 0x7c00, is copied to register %si. The destination address, 0x600, to register %di. The number of words to copy, 256 (the program’s size = 512 bytes), is copied to register %cx. Next, the rep instruction repeats the instruction that follows, that is, movsw, the number of times dictated by the %cx register. The movsw instruction copies the word pointed to by %si to the address pointed to by %di. This is repeated another 255 times. On each repetition, both the source and destination registers, %si and %di, are incremented by one. Thus, upon completion of the 256-word (512-byte) copy, %di has the value 0x600`512`= `0x800`, and `%si` has the value `0x7c00`512= 0x7e00; we have thus completed the code relocation. Since the last update of this document, the copy instructions have changed in the code, so instead of the movsb and stosb, movsw and stosw have been introduced, which copy 2 bytes(1 word) in one iteration.

Next, the destination register %di is copied to %bp. %bp gets the value 0x800. The value 8 is copied to %cl in preparation for a new string operation (like our previous movsw). Now, stosw is executed 8 times. This instruction copies a 0 value to the address pointed to by the destination register (%di, which is 0x800), and increments it. This is repeated another 7 times, so %di ends up with value 0x810. Effectively, this clears the address range 0x800-0x80f. This range is used as a (fake) partition table for writing the MBR back to disk. Finally, the sector field for the CHS addressing of this fake partition is given the value 1 and a jump is made to the main function from the relocated code. Note that until this jump to the relocated code, any reference to an absolute address was avoided.

The following code block tests whether the drive number provided by the BIOS should be used, or the one stored in boot0.

This code tests the SETDRV bit (0x20) in the flags variable. Recall that register %bp points to address location 0x800, so the test is done to the flags variable at address 0x800-69= 0x7bb. This is an example of the type of modifications that can be done to boot0. The SETDRV flag is not set by default, but it can be set in the Makefile. When set, the drive number stored in the MBR is used instead of the one provided by the BIOS. We assume the defaults, and that the BIOS provided a valid drive number, so we jump to save_curdrive.

The next block saves the drive number provided by the BIOS, and calls putn to print a new line on the screen.

Note that we assume TEST is not defined, so the conditional code in it is not assembled and will not appear in our executable boot0.

Our next block implements the actual scanning of the partition table. It prints to the screen the partition type for each of the four entries in the partition table. It compares each type with a list of well-known operating system file systems. Examples of recognized partition types are NTFS (Windows®, ID 0x7), ext2fs (Linux®, ID 0x83), and, of course, ffs/ufs2 (FreeBSD, ID 0xa5). The implementation is fairly simple.

It is important to note that the active flag for each entry is cleared, so after the scanning, no partition entry is active in our memory copy of boot0. Later, the active flag will be set for the selected partition. This ensures that only one active partition exists if the user chooses to write the changes back to disk.

The next block tests for other drives. At startup, the BIOS writes the number of drives present in the computer to address 0x475. If there are any other drives present, boot0 prints the current drive to screen. The user may command boot0 to scan partitions on another drive later.

We make the assumption that a single drive is present, so the jump to print_drive is not performed. We also assume nothing strange happened, so we jump to print_prompt.

This next block just prints out a prompt followed by the default option:

Finally, a jump is performed to start_input, where the BIOS services are used to start a timer and for reading user input from the keyboard; if the timer expires, the default option will be selected:

An interrupt is requested with number 0x1a and argument 0 in register %ah. The BIOS has a predefined set of services, requested by applications as software-generated interrupts through the int instruction and receiving arguments in registers (in this case, %ah). Here, particularly, we are requesting the number of clock ticks since last midnight; this value is computed by the BIOS through the RTC (Real Time Clock). This clock can be programmed to work at frequencies ranging from 2 Hz to 8192 Hz. The BIOS sets it to 18.2 Hz at startup. When the request is satisfied, a 32-bit result is returned by the BIOS in registers %cx and %dx (lower bytes in %dx). This result (the %dx part) is copied to register %di, and the value of the TICKS variable is added to %di. This variable resides in boot0 at offset _TICKS (a negative value) from register %bp (which, recall, points to 0x800). The default value of this variable is 0xb6 (182 in decimal). Now, the idea is that boot0 constantly requests the time from the BIOS, and when the value returned in register %dx is greater than the value stored in %di, the time is up and the default selection will be made. Since the RTC ticks 18.2 times per second, this condition will be met after 10 seconds (this default behavior can be changed in the Makefile). Until this time has passed, boot0 continually asks the BIOS for any user input; this is done through int 0x16, argument 1 in %ah.

Whether a key was pressed or the time expired, subsequent code validates the selection. Based on the selection, the register %si is set to point to the appropriate partition entry in the partition table. This new selection overrides the previous default one. Indeed, it becomes the new default. Finally, the ACTIVE flag of the selected partition is set. If it was enabled at compile time, the in-memory version of boot0 with these modified values is written back to the MBR on disk. We leave the details of this implementation to the reader.

We now end our study with the last code block from the boot0 program:

Recall that %si points to the selected partition entry. This entry tells us where the partition begins on disk. We assume, of course, that the partition selected is actually a FreeBSD slice.

The transfer buffer is set to 0x7c00 (register %bx), and a read for the first sector of the FreeBSD slice is requested by calling intx13. We assume that everything went okay, so a jump to beep is not performed. In particular, the new sector read must end with the magic sequence 0xaa55. Finally, the value at %si (the pointer to the selected partition table) is preserved for use by the next stage, and a jump is performed to address 0x7c00, where execution of our next stage (the just-read block) is started.

So far we have gone through the following sequence:

boot1 is the next step in the boot-loading sequence. It is the first of three boot stages. Note that we have been dealing exclusively with disk sectors. Indeed, the BIOS loads the absolute first sector, while boot0 loads the first sector of the FreeBSD slice. Both loads are to address 0x7c00. We can conceptually think of these disk sectors as containing the files boot0 and boot1, respectively, but in reality this is not entirely true for boot1. Strictly speaking, unlike boot0, boot1 is not part of the boot blocks ^[3]. Instead, a single, full-blown file, boot (/boot/boot), is what ultimately is written to disk. This file is a combination of boot1, boot2 and the Boot Extender (or BTX). This single file is greater in size than a single sector (greater than 512 bytes). Fortunately, boot1 occupies exactly the first 512 bytes of this single file, so when boot0 loads the first sector of the FreeBSD slice (512 bytes), it is actually loading boot1 and transferring control to it.

The main task of boot1 is to load the next boot stage. This next stage is somewhat more complex. It is composed of a server called the "Boot Extender", or BTX, and a client, called boot2. As we will see, the last boot stage, loader, is also a client of the BTX server.

Let us now look in detail at what exactly is done by boot1, starting like we did for boot0, at its entry point:

The entry point at start simply jumps past a special data area to the label main, which in turn looks like this:

Just like boot0, this code relocates boot1, this time to memory address 0x700. However, unlike boot0, it does not jump there. boot1 is linked to execute at address 0x7c00, effectively where it was loaded in the first place. The reason for this relocation will be discussed shortly.

Next comes a loop that looks for the FreeBSD slice. Although boot0 loaded boot1 from the FreeBSD slice, no information was passed to it about this ^[4], so boot1 must rescan the partition table to find where the FreeBSD slice starts. Therefore it rereads the MBR:

In the code above, register %dl maintains information about the boot device. This is passed on by the BIOS and preserved by the MBR. Numbers 0x80 and greater tells us that we are dealing with a hard drive, so a call is made to nread, where the MBR is read. Arguments to nread are passed through %si and %dh. The memory address at label part4 is copied to %si. This memory address holds a "fake partition" to be used by nread. The following is the data in the fake partition:

In particular, the LBA for this fake partition is hardcoded to zero. This is used as an argument to the BIOS for reading absolute sector one from the hard drive. Alternatively, CHS addressing could be used. In this case, the fake partition holds cylinder 0, head 0 and sector 1, which is equivalent to absolute sector one.

Let us now proceed to take a look at nread:

Recall that %si points to the fake partition. The word ^[5] at offset 0x8 is copied to register %ax and word at offset 0xa to %cx. They are interpreted by the BIOS as the lower 4-byte value denoting the LBA to be read (the upper four bytes are assumed to be zero). Register %bx holds the memory address where the MBR will be loaded. The instruction pushing %cs onto the stack is very interesting. In this context, it accomplishes nothing. However, as we will see shortly, boot2, in conjunction with the BTX server, also uses xread.1. This mechanism will be discussed in the next section.

The code at xread.1 further calls the read function, which actually calls the BIOS asking for the disk sector:

Note the long return instruction at the end of this block. This instruction pops out the %cs register pushed by nread, and returns. Finally, nread also returns.

With the MBR loaded to memory, the actual loop for searching the FreeBSD slice begins:

If a FreeBSD slice is identified, execution continues at main.5. Note that when a FreeBSD slice is found %si points to the appropriate entry in the partition table, and %dh holds the partition number. We assume that a FreeBSD slice is found, so we continue execution at main.5:

Recall that at this point, register %si points to the FreeBSD slice entry in the MBR partition table, so a call to nread will effectively read sectors at the beginning of this partition. The argument passed on register %dh tells nread to read 16 disk sectors. Recall that the first 512 bytes, or the first sector of the FreeBSD slice, coincides with the boot1 program. Also recall that the file written to the beginning of the FreeBSD slice is not /boot/boot1, but /boot/boot. Let us look at the size of these files in the filesystem:

Both boot0 and boot1 are 512 bytes each, so they fit exactly in one disk sector. boot2 is much bigger, holding both the BTX server and the boot2 client. Finally, a file called simply boot is 512 bytes larger than boot2. This file is a concatenation of boot1 and boot2. As already noted, boot0 is the file written to the absolute first disk sector (the MBR), and boot is the file written to the first sector of the FreeBSD slice; boot1 and boot2 are not written to disk. The command used to concatenate boot1 and boot2 into a single boot is merely cat boot1 boot2 > boot.

So boot1 occupies exactly the first 512 bytes of boot and, because boot is written to the first sector of the FreeBSD slice, boot1 fits exactly in this first sector. When nread reads the first 16 sectors of the FreeBSD slice, it effectively reads the entire boot file ^[6]. We will see more details about how boot is formed from boot1 and boot2 in the next section.

Recall that nread uses memory address 0x8c00 as the transfer buffer to hold the sectors read. This address is conveniently chosen. Indeed, because boot1 belongs to the first 512 bytes, it ends up in the address range 0x8c00-0x8dff. The 512 bytes that follows (range 0x8e00-0x8fff) is used to store the bsdlabel ^[7].

Starting at address 0x9000 is the beginning of the BTX server, and immediately following is the boot2 client. The BTX server acts as a kernel, and executes in protected mode in the most privileged level. In contrast, the BTX clients (boot2, for example), execute in user mode. We will see how this is accomplished in the next section. The code after the call to nread locates the beginning of boot2 in the memory buffer, and copies it to memory address 0xc000. This is because the BTX server arranges boot2 to execute in a segment starting at 0xa000. We explore this in detail in the following section.

The last code block of boot1 enables access to memory above 1MB ^[8] and concludes with a jump to the starting point of the BTX server:

Note that right before the jump, interrupts are enabled.

Next in our boot sequence is the BTX Server. Let us quickly remember how we got here:

Before studying the BTX Server in detail, let us further review how the single, all-in-one boot file is created. The way boot is built is defined in its Makefile (stand/i386/boot2/Makefile). Let us look at the rule that creates the boot file:

This tells us that boot1 and boot2 are needed, and the rule simply concatenates them to produce a single file called boot. The rules for creating boot1 are also quite simple:

To apply the rule for creating boot1, boot1.out must be resolved. This, in turn, depends on the existence of boot1.o. This last file is simply the result of assembling our familiar boot1.S, without linking. Now, the rule for creating boot1.out is applied. This tells us that boot1.o should be linked with start as its entry point, and starting at address 0x7c00. Finally, boot1 is created from boot1.out applying the appropriate rule. This rule is the objcopy command applied to boot1.out. Note the flags passed to objcopy: -S tells it to strip all relocation and symbolic information; -O binary indicates the output format, that is, a simple, unformatted binary file.

Having boot1, let us take a look at how boot2 is constructed:

The mechanism for building boot2 is far more elaborate. Let us point out the most relevant facts. The dependency list is as follows:

Note that initially there is no header file boot2.h, but its creation depends on boot1.out, which we already have. The rule for its creation is a bit terse, but the important thing is that the output, boot2.h, is something like this:

Recall that boot1 was relocated (i.e., copied from 0x7c00 to 0x700). This relocation will now make sense, because as we will see, the BTX server reclaims some memory, including the space where boot1 was originally loaded. However, the BTX server needs access to boot1's xread function; this function, according to the output of boot2.h, is at location 0x725. Indeed, the BTX server uses the xread function from boot1's relocated code. This function is now accessible from within the boot2 client.

The next rule directs the linker to link various files (ashldi3.o, boot2.o and sio.o). Note that the output file, boot2.out, is linked to execute at address 0x2000 (${ORG2}). Recall that boot2 will be executed in user mode, within a special user segment set up by the BTX server. This segment starts at 0xa000. Also, remember that the boot2 portion of boot was copied to address 0xc000, that is, offset 0x2000 from the start of the user segment, so boot2 will work properly when we transfer control to it. Next, boot2.bin is created from boot2.out by stripping its symbols and format information; boot2.bin is a raw binary. Now, note that a file boot2.ldr is created as a 512-byte file full of zeros. This space is reserved for the bsdlabel.

Now that we have files boot1, boot2.bin and boot2.ldr, only the BTX server is missing before creating the all-in-one boot file. The BTX server is located in stand/i386/btx/btx; it has its own Makefile with its own set of rules for building. The important thing to notice is that it is also compiled as a raw binary, and that it is linked to execute at address 0x9000. The details can be found in stand/i386/btx/btx/Makefile.

Having the files that comprise the boot program, the final step is to merge them. This is done by a special program called btxld (source located in /usr/src/usr.sbin/btxld). Some arguments to this program include the name of the output file (boot), its entry point (0x2000) and its file format (raw binary). The various files are finally merged by this utility into the file boot, which consists of boot1, boot2, the bsdlabel and the BTX server. This file, which takes exactly 16 sectors, or 8192 bytes, is what is actually written to the beginning of the FreeBSD slice during installation. Let us now proceed to study the BTX server program.

The BTX server prepares a simple environment and switches from 16-bit real mode to 32-bit protected mode, right before passing control to the client. This includes initializing and updating the following data structures:

Let us now start studying the actual implementation. Recall that boot1 made a jump to address 0x9010, the BTX server’s entry point. Before studying program execution there, note that the BTX server has a special header at address range 0x9000-0x900f, right before its entry point. This header is defined as follows:

Note the first two bytes are 0xeb and 0xe. In the IA-32 architecture, these two bytes are interpreted as a relative jump past the header into the entry point, so in theory, boot1 could jump here (address 0x9000) instead of address 0x9010. Note that the last field in the BTX header is a pointer to the client’s (boot2) entry pointb2. This field is patched at link time.

Immediately following the header is the BTX server’s entry point:

This code disables interrupts, sets up a working stack (starting at address 0x1800) and clears the flags in the EFLAGS register. Note that the popfl instruction pops out a doubleword (4 bytes) from the stack and places it in the EFLAGS register. As the value actually popped is 2, the EFLAGS register is effectively cleared (IA-32 requires that bit 2 of the EFLAGS register always be 1).

Our next code block clears (sets to 0) the memory range 0x5e00-0x8fff. This range is where the various data structures will be created:

Recall that boot1 was originally loaded to address 0x7c00, so, with this memory initialization, that copy effectively disappeared. However, also recall that boot1 was relocated to 0x700, so that copy is still in memory, and the BTX server will make use of it.

Next, the real-mode IVT (Interrupt Vector Table is updated. The IVT is an array of segment/offset pairs for exception and interrupt handlers. The BIOS normally maps hardware interrupts to interrupt vectors 0x8 to 0xf and 0x70 to 0x77 but, as will be seen, the 8259A Programmable Interrupt Controller, the chip controlling the actual mapping of hardware interrupts to interrupt vectors, is programmed to remap these interrupt vectors from 0x8-0xf to 0x20-0x27 and from 0x70-0x77 to 0x28-0x2f. Thus, interrupt handlers are provided for interrupt vectors 0x20-0x2f. The reason the BIOS-provided handlers are not used directly is because they work in 16-bit real mode, but not 32-bit protected mode. Processor mode will be switched to 32-bit protected mode shortly. However, the BTX server sets up a mechanism to effectively use the handlers provided by the BIOS:

The next block creates the IDT (Interrupt Descriptor Table). The IDT is analogous, in protected mode, to the IVT in real mode. That is, the IDT describes the various exception and interrupt handlers used when the processor is executing in protected mode. In essence, it also consists of an array of segment/offset pairs, although the structure is somewhat more complex, because segments in protected mode are different than in real mode, and various protection mechanisms apply:

Each entry in the IDT is 8 bytes long. Besides the segment/offset information, they also describe the segment type, privilege level, and whether the segment is present in memory or not. The construction is such that interrupt vectors from 0 to 0xf (exceptions) are handled by function intx00; vector 0x10 (also an exception) is handled by intx10; hardware interrupts, which are later configured to start at interrupt vector 0x20 all the way to interrupt vector 0x2f, are handled by function intx20. Lastly, interrupt vector 0x30, which is used for system calls, is handled by intx30, and vectors 0x31 and 0x32 are handled by intx31. It must be noted that only descriptors for interrupt vectors 0x30, 0x31 and 0x32 are given privilege level 3, the same privilege level as the boot2 client, which means the client can execute a software-generated interrupt to this vectors through the int instruction without failing (this is the way boot2 use the services provided by the BTX server). Also, note that only software-generated interrupts are protected from code executing in lesser privilege levels. Hardware-generated interrupts and processor-generated exceptions are always handled adequately, regardless of the actual privileges involved.

The next step is to initialize the TSS (Task-State Segment). The TSS is a hardware feature that helps the operating system or executive software implement multitasking functionality through process abstraction. The IA-32 architecture demands the creation and use of at least one TSS if multitasking facilities are used or different privilege levels are defined. Since the boot2 client is executed in privilege level 3, but the BTX server runs in privilege level 0, a TSS must be defined:

Note that a value is given for the Privilege Level 0 stack pointer and stack segment in the TSS. This is needed because, if an interrupt or exception is received while executing boot2 in Privilege Level 3, a change to Privilege Level 0 is automatically performed by the processor, so a new working stack is needed. Finally, the I/O Map Base Address field of the TSS is given a value, which is a 16-bit offset from the beginning of the TSS to the I/O Permission Bitmap and the Interrupt Redirection Bitmap.

After the IDT and TSS are created, the processor is ready to switch to protected mode. This is done in the next block:

First, a call is made to setpic to program the 8259A PIC (Programmable Interrupt Controller). This chip is connected to multiple hardware interrupt sources. Upon receiving an interrupt from a device, it signals the processor with the appropriate interrupt vector. This can be customized so that specific interrupts are associated with specific interrupt vectors, as explained before. Next, the IDTR (Interrupt Descriptor Table Register) and GDTR (Global Descriptor Table Register) are loaded with the instructions lidt and lgdt, respectively. These registers are loaded with the base address and limit address for the IDT and GDT. The following three instructions set the Protection Enable (PE) bit of the %cr0 register. This effectively switches the processor to 32-bit protected mode. Next, a long jump is made to init.8 using segment selector SEL_SCODE, which selects the Supervisor Code Segment. The processor is effectively executing in CPL 0, the most privileged level, after this jump. Finally, the Supervisor Data Segment is selected for the stack by assigning the segment selector SEL_SDATA to the %ss register. This data segment also has a privilege level of 0.

Our last code block is responsible for loading the TR (Task Register) with the segment selector for the TSS we created earlier, and setting the User Mode environment before passing execution control to the boot2 client.

Note that the client’s environment include a stack segment selector and stack pointer (registers %ss and %esp). Indeed, once the TR is loaded with the appropriate stack segment selector (instruction ltr), the stack pointer is calculated and pushed onto the stack along with the stack’s segment selector. Next, the value 0x202 is pushed onto the stack; it is the value that the EFLAGS will get when control is passed to the client. Also, the User Mode code segment selector and the client’s entry point are pushed. Recall that this entry point is patched in the BTX header at link time. Finally, segment selectors (stored in register %ecx) for the segment registers %gs, %fs, %ds and %es are pushed onto the stack, along with the value at %edx (0xa000). Keep in mind the various values that have been pushed onto the stack (they will be popped out shortly). Next, values for the remaining general purpose registers are also pushed onto the stack (note the loop that pushes the value 0 seven times). Now, values will be started to be popped out of the stack. First, the popa instruction pops out of the stack the latest seven values pushed. They are stored in the general purpose registers in order %edi, %esi, %ebp, %ebx, %edx, %ecx, %eax. Then, the various segment selectors pushed are popped into the various segment registers. Five values still remain on the stack. They are popped when the iret instruction is executed. This instruction first pops the value that was pushed from the BTX header. This value is a pointer to boot2's entry point. It is placed in the register %eip, the instruction pointer register. Next, the segment selector for the User Code Segment is popped and copied to register %cs. Remember that this segment’s privilege level is 3, the least privileged level. This means that we must provide values for the stack of this privilege level. This is why the processor, besides further popping the value for the EFLAGS register, does two more pops out of the stack. These values go to the stack pointer (%esp) and the stack segment (%ss). Now, execution continues at boot0's entry point.

It is important to note how the User Code Segment is defined. This segment’s base address is set to 0xa000. This means that code memory addresses are relative to address 0xa000; if code being executed is fetched from address 0x2000, the actual memory addressed is 0xa000+0x2000=0xc000.

boot2 defines an important structure, struct bootinfo. This structure is initialized by boot2 and passed to the loader, and then further to the kernel. Some nodes of this structures are set by boot2, the rest by the loader. This structure, among other information, contains the kernel filename, BIOS harddisk geometry, BIOS drive number for boot device, physical memory available, envp pointer etc. The definition for it is:

boot2 enters into an infinite loop waiting for user input, then calls load(). If the user does not press anything, the loop breaks by a timeout, so load() will load the default file (/boot/loader). Functions ino_t lookup(char *filename) and int xfsread(ino_t inode, void *buf, size_t nbyte) are used to read the content of a file into memory. /boot/loader is an ELF binary, but where the ELF header is prepended with a.out's struct exec structure. load() scans the loader’s ELF header, loading the content of /boot/loader into memory, and passing the execution to the loader’s entry:

loader is a BTX client as well. I will not describe it here in detail, there is a comprehensive man page written by Mike Smith, loader(8). The underlying mechanisms and BTX were discussed above.

The main task for the loader is to boot the kernel. When the kernel is loaded into memory, it is being called by the loader:

Let us take a look at the command that links the kernel. This will help identify the exact location where the loader passes execution to the kernel. This location is the kernel’s actual entry point. This command is now excluded from sys/conf/Makefile.i386. The content that interests us can be found in /usr/obj/usr/src/i386.i386/sys/GENERIC/.

A few interesting things can be seen here. First, the kernel is an ELF dynamically linked binary, but the dynamic linker for kernel is /red/herring, which is definitely a bogus file. Second, taking a look at the file sys/conf/ldscript.i386 gives an idea about what ld options are used when compiling a kernel. Reading through the first few lines, the string

says that a kernel’s entry point is the symbol btext. This symbol is defined in locore.s:

First, the register EFLAGS is set to a predefined value of 0x00000002. Then all the segment registers are initialized:

btext calls the routines recover_bootinfo(), identify_cpu(), which are also defined in locore.s. Here is a description of what they do:

The next steps are enabling VME, if the CPU supports it:

Then, enabling paging:

The next three lines of code are because the paging was set, so the jump is needed to continue the execution in virtualized address space:

The function init386() is called with a pointer to the first free physical page, after that mi_startup(). init386 is an architecture dependent initialization function, and mi_startup() is an architecture independent one (the 'mi_' prefix stands for Machine Independent). The kernel never returns from mi_startup(), and by calling it, the kernel finishes booting:

A mutex is simply a lock used to guarantee mutual exclusion. Specifically, a mutex may only be owned by one entity at a time. If another entity wishes to obtain a mutex that is already owned, it must wait until the mutex is released. In the FreeBSD kernel, mutexes are owned by processes.

Mutexes may be recursively acquired, but they are intended to be held for a short period of time. Specifically, one may not sleep while holding a mutex. If you need to hold a lock across a sleep, use a lockmgr(9) lock.

Each mutex has several properties of interest:

These locks provide basic reader-writer type functionality and may be held by a sleeping process. Currently they are backed by lockmgr(9).

An atomically protected variable is a special variable that is not protected by an explicit lock. Instead, all data accesses to the variables use special atomic operations as described in atomic(9). Very few variables are treated this way, although other synchronization primitives such as mutexes are implemented with atomically protected variables.

Kobj works by generating descriptions of methods. Each description holds a unique id as well as a default function. The description’s address is used to uniquely identify the method within a class' method table.

A class is built by creating a method table associating one or more functions with method descriptions. Before use the class is compiled. The compilation allocates a cache and associates it with the class. A unique id is assigned to each method description within the method table of the class if not already done so by another referencing class compilation. For every method to be used a function is generated by script to qualify arguments and automatically reference the method description for a lookup. The generated function looks up the method by using the unique id associated with the method description as a hash into the cache associated with the object’s class. If the method is not cached the generated function proceeds to use the class' table to find the method. If the method is found then the associated function within the class is used; otherwise, the default function associated with the method description is used.

These indirections can be visualized as the following:

Jail consists of two realms: the userland program, jail(8), and the code implemented within the kernel: the jail(2) system call and associated restrictions. I will be discussing the userland program and then how jail is implemented within the kernel.

Throughout the kernel there are access restrictions relating to jailed processes. Usually, these restrictions only check whether the process is jailed, and if so, returns an error. For example:

SYSINIT relies on the ability of the linker to take static data declared at multiple locations throughout a program’s source and group it together as a single contiguous chunk of data. This linker technique is called a "linker set". SYSINIT uses two linker sets to maintain two data sets containing each consumer’s call order, function, and a pointer to the data to pass to that function.

SYSINIT uses two priorities when ordering the functions for execution. The first priority is a subsystem ID giving an overall order for SYSINIT’s dispatch of functions. Current predeclared ID’s are in <sys/kernel.h> in the enum list sysinit_sub_id. The second priority used is an element order within the subsystem. Current predeclared subsystem element orders are in <sys/kernel.h> in the enum list sysinit_elem_order.

There are currently two uses for SYSINIT. Function dispatch at system startup and kernel module loads, and function dispatch at system shutdown and kernel module unload. Kernel subsystems often use system startup SYSINIT’s to initialize data structures, for example the process scheduling subsystem uses a SYSINIT to initialize the run queue data structure. Device drivers should avoid using SYSINIT() directly. Instead drivers for real devices that are part of a bus structure should use DRIVER_MODULE() to provide a function that detects the device and, if it is present, initializes the device. It will do a few things specific to devices and then call SYSINIT() itself. For pseudo-devices, which are not part of a bus structure, use DEV_MODULE().

This documentation was developed for the FreeBSD Project by Chris Costello at Safeport Network Services and Network Associates Laboratories, the Security Research Division of Network Associates, Inc. under DARPA/SPAWAR contract N66001-01-C-8035 ("CBOSS"), as part of the DARPA CHATS research program.

Redistribution and use in source (SGML DocBook) and 'compiled' forms (SGML, HTML, PDF, PostScript, RTF and so forth) with or without modification, are permitted provided that the following conditions are met:

FreeBSD includes experimental support for several mandatory access control policies, as well as a framework for kernel security extensibility, the TrustedBSD MAC Framework. The MAC Framework is a pluggable access control framework, permitting new security policies to be easily linked into the kernel, loaded at boot, or loaded dynamically at run-time. The framework provides a variety of features to make it easier to implement new security policies, including the ability to easily tag security labels (such as confidentiality information) onto system objects.

This chapter introduces the MAC policy framework and provides documentation for a sample MAC policy module.

The TrustedBSD MAC framework provides a mechanism to allow the compile-time or run-time extension of the kernel access control model. New system policies may be implemented as kernel modules and linked to the kernel; if multiple policy modules are present, their results will be composed. The MAC Framework provides a variety of access control infrastructure services to assist policy writers, including support for transient and persistent policy-agnostic object security labels. This support is currently considered experimental.

This chapter provides information appropriate for developers of policy modules, as well as potential consumers of MAC-enabled environments, to learn about how the MAC Framework supports access control extension of the kernel.

Mandatory Access Control (MAC), refers to a set of access control policies that are mandatorily enforced on users by the operating system. MAC policies may be contrasted with Discretionary Access Control (DAC) protections, by which non-administrative users may (at their discretion) protect objects. In traditional UNIX systems, DAC protections include file permissions and access control lists; MAC protections include process controls preventing inter-user debugging and firewalls. A variety of MAC policies have been formulated by operating system designers and security researches, including the Multi-Level Security (MLS) confidentiality policy, the Biba integrity policy, Role-Based Access Control (RBAC), Domain and Type Enforcement (DTE), and Type Enforcement (TE). Each model bases decisions on a variety of factors, including user identity, role, and security clearance, as well as security labels on objects representing concepts such as data sensitivity and integrity.

The TrustedBSD MAC Framework is capable of supporting policy modules that implement all of these policies, as well as a broad class of system hardening policies, which may use existing security attributes, such as user and group IDs, as well as extended attributes on files, and other system properties. In addition, despite the name, the MAC Framework can also be used to implement purely discretionary policies, as policy modules are given substantial flexibility in how they authorize protections.

The TrustedBSD MAC Framework permits kernel modules to extend the operating system security policy, as well as providing infrastructure functionality required by many access control modules. If multiple policies are simultaneously loaded, the MAC Framework will usefully (for some definition of useful) compose the results of the policies.

Security policies are either linked directly into the kernel, or compiled into loadable kernel modules that may be loaded at boot, or dynamically using the module loading system calls at runtime. Policy modules interact with the system through a set of declared entry points, providing access to a stream of system events and permitting the policy to influence access control decisions. Each policy contains a number of elements: