Chapter 6: Basic Technique Limited Direct Execution

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• Earlier prerequisite concepts leading into Chapter 6: Basic Technique Limited Direct Execution.

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• Raw source file: 029-6-1-basic-technique-limited-direct-execution.md
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• Raw source file: 031-6-3-problem-2-switching-between-processes.md
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Basic Technique Limited Direct Execution

6.1 Basic Technique: Limited Direct Execution

6 Mechanism: Limited Direct Execution

In order to virtualize the CPU, the operating system needs to somehow share the physical CPU among many jobs running seemingly at the same time. The basic idea is simple: run one process for a little while, then run another one, and so forth. Bytime sharingthe CPU in this manner, virtualization is achieved.

There are a few challenges, however, in building such virtualization machinery. The first is performance: how can we implement virtualization without adding excessive overhead to the system? The second is control: how can we run processes efficiently while retaining control over the CPU? Control is particularly important to the OS, as it is in charge of resources; without control, a process could simply run forever and take over the machine, or access information that it should not be allowed to access. Obtaining high performance while maintaining control is thus one of the central challenges in building an operating system.

THECRUX:

HOWTOEFFICIENTLYVIRTUALIZETHECPU WITHCONTROL

The OS must virtualize the CPU in an efficient manner while retaining control over the system. To do so, both hardware and operating-system support will be required. The OS will often use a judicious bit of hardware support in order to accomplish its work effectively.

To make a program run as fast as one might expect, not surprisingly

OS developers came up with a technique, which we calllimited direct execution. The "direct execution" part of the idea is simple: just run the program directly on the CPU. Thus, when the OS wishes to start a program running, it creates a process entry for it in a process list, allocates some memory for it, loads the program code into memory (from disk), locates its entry point (i.e., themain()routine or something similar), jumps 1

OS Program

Create entry for process list

Allocate memory for program

Load program into memory

Set up stack with argc/argv

Clear registers

Executecallmain()

Run main()

Executereturnfrom main

Free memory of process

Remove from process list

Figure 6.1:Direct Execution Protocol (Without Limits) to it, and starts running the user's code. Figure 6.1 shows this basic direct execution protocol (without any limits, yet), using a normal call and return to jump to the program'smain()and later to get back into the kernel.

Sounds simple, no? But this approach gives rise to a few problems in our quest to virtualize the CPU. The first is simple: if we just run a program, how can the OS make sure the program doesn't do anything that we don't want it to do, while still running it efficiently? The second:

when we are running a process, how does the operating system stop it from running and switch to another process, thus implementing thetime sharingwe require to virtualize the CPU?

In answering these questions below, we'll get a much better sense of what is needed to virtualize the CPU. In developing these techniques, we'll also see where the "limited" part of the name arises from; without limits on running programs, the OS wouldn't be in control of anything and thus would be "just a library" -- a very sad state of affairs for an aspiring operating system!

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Problem 1 Restricted Operations

6.2 Problem #1: Restricted Operations

Direct execution has the obvious advantage of being fast; the program runs natively on the hardware CPU and thus executes as quickly as one would expect. But running on the CPU introduces a problem: what if the process wishes to perform some kind of restricted operation, such as issuing an I/O request to a disk, or gaining access to more system resources such as CPU or memory?

THECRUX: HOWTOPERFORMRESTRICTEDOPERATIONS

A process must be able to perform I/O and some other restricted operations, but without giving the process complete control over the system.

How can the OS and hardware work together to do so?

TIP: USEPROTECTEDCONTROLTRANSFER

The hardware assists the OS by providing different modes of execution.

Inuser mode, applications do not have full access to hardware resources.

Inkernel mode, the OS has access to the full resources of the machine.

Special instructions totrapinto the kernel andreturn-from-trapback to user-mode programs are also provided, as well instructions that allow the

OS to tell the hardware where thetrap tableresides in memory.

One approach would simply be to let any process do whatever it wants in terms of I/O and other related operations. However, doing so would prevent the construction of many kinds of systems that are desirable. For

example, if we wish to build a file system that checks permissions before granting access to a file, we can't simply let any user process issue I/Os to the disk; if we did, a process could simply read or write the entire disk and thus all protections would be lost.

Thus, the approach we take is to introduce a new processor mode, known asuser mode; code that runs in user mode is restricted in what it can do. For example, when running in user mode, a process can't issue

I/O requests; doing so would result in the processor raising an exception; the OS would then likely kill the process.

In contrast to user mode iskernel mode, which the operating system (or kernel) runs in. In this mode, code that runs can do what it likes, including privileged operations such as issuing I/O requests and executing all types of restricted instructions.

We are still left with a challenge, however: what should a user process do when it wishes to perform some kind of privileged operation, such as reading from disk? To enable this, virtually all modern hardware provides the ability for user programs to perform a system call.

Pioneered on ancient machines such as the Atlas [K+61,L78], system calls allow the kernel to carefully expose certain key pieces of functionality to user programs, such as accessing the file system, creating and destroying processes, communicating with other processes, and allocating more memory. Most operating systems provide a few hundred calls (see the

POSIX standard for details [P10]); early Unix systems exposed a more concise subset of around twenty calls.

To execute a system call, a program must execute a specialtrapinstruction. This instruction simultaneously jumps into the kernel and raises the privilege level to kernel mode; once in the kernel, the system can now perform whatever privileged operations are needed (if allowed), and thus do the required work for the calling process. When finished, the OS calls a specialreturn-from-trapinstruction, which, as you might expect, returns into the calling user program while simultaneously reducing the privilege level back to user mode.

The hardware needs to be a bit careful when executing a trap, in that it must make sure to save enough of the caller's registers in order to be able

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Problem 2 Switching Between Processes

6.3 Problem #2: Switching Between Processes

ASIDE: WHYSYSTEMCALLSLOOKLIKEPROCEDURECALLS

You may wonder why a call to a system call, such asopen()orread(), looks exactly like a typical procedure call in C; that is, if it looks just like a procedure call, how does the system know it's a system call, and do all the right stuff? The simple reason: itis a procedure call, but hidden inside that procedure call is the famous trap instruction. More specifically, when you callopen()(for example), you are executing a procedure call into the C library. Therein, whether foropen()or any of the other system calls provided, the library uses an agreed-upon calling convention with the kernel to put the arguments to open in well-known locations (e.g., on the stack, or in specific registers), puts the system-call number into a well-known location as well (again, onto the stack or a register), and then executes the aforementioned trap instruction. The code in the library after the trap unpacks return values and returns control to the program that issued the system call. Thus, the parts of the C library that make system calls are hand-coded in assembly, as they need to carefully follow convention in order to process arguments and return values correctly, as well as execute the hardware-specific trap instruction. And now you know why you personally don't have to write assembly code to trap into an OS; somebody has already written that assembly for you.

to return correctly when the OS issues the return-from-trap instruction.

On x86, for example, the processor will push the program counter, flags, and a few other registers onto a per-processkernel stack; the return-fromtrap will pop these values off the stack and resume execution of the usermode program (see the Intel systems manuals [I11] for details). Other hardware systems use different conventions, but the basic concepts are similar across platforms.

There is one important detail left out of this discussion: how does the trap know which code to run inside the OS? Clearly, the calling process can't specify an address to jump to (as you would when making a procedure call); doing so would allow programs to jump anywhere into the kernel which clearly is a bad idea (imagine jumping into code to access a file, but just after a permission check; in fact, it is likely such an ability would enable a wily programmer to get the kernel to run arbitrary code sequences [S07]). Thus the kernel must carefully control what code executes upon a trap.

The kernel does so by setting up atrap tableat boot time. When the machine boots up, it does so in privileged (kernel) mode, and thus is free to configure machine hardware as need be. One of the first things the OS thus does is to tell the hardware what code to run when certain exceptional events occur. For example, what code should run when a harddisk interrupt takes place, when a keyboard interrupt occurs, or when a program makes a system call? The OS informs the hardware of the locations of these trap handlers, usually with some kind of special ininitialize trap table remember address of...

syscall handler

OS @ run Hardware Program

(kernel mode) (user mode)

Create entry for process list

Allocate memory for program

Load program into memory

Setup user stack with argv

Fill kernel stack with reg/PC

return-from-trap

restore regs from kernel stack move to user mode jump to main

Run main()
...

Call system call trapinto OS save regs to kernel stack move to kernel mode jump to trap handler

Handle trap

Do work of syscall

return-from-trap

restore regs from kernel stack move to user mode jump to PC after trap

...
return from main
trap(viaexit())

Free memory of process

Remove from process list

Figure 6.2:Limited Direct Execution Protocol struction. Once the hardware is informed, it remembers the location of these handlers until the machine is next rebooted, and thus the hardware knows what to do (i.e., what code to jump to) when system calls and other exceptional events take place.

One last aside: being able to execute the instruction to tell the hardware where the trap tables are is a very powerful capability. Thus, as you might have guessed, it is also aprivilegedoperation. If you try to execute this instruction in user mode, the hardware won't let you, and you can probably guess what will happen (hint: adios, offending program).

Point to ponder: what horrible things could you do to a system if you could install your own trap table? Could you take over the machine?

The timeline (with time increasing downward, in Figure 6.2) summarizes the protocol. We assume each process has a kernel stack where registers (including general purpose registers and the program counter) are saved to and restored from (by the hardware) when transitioning into and out of the kernel.

There are two phases in the LDE protocol. In the first (at boot time), the kernel initializes the trap table, and the CPU remembers its location for subsequent use. The kernel does so via a privileged instruction (all privileged instructions are highlighted in bold).

In the second (when running a process), the kernel sets up a few things (e.g., allocating a node on the process list, allocating memory) before using a return-from-trap instruction to start the execution of the process;

this switches the CPU to user mode and begins running the process.

When the process wishes to issue a system call, it traps back into the OS, which handles it and once again returns control via a return-from-trap to the process. The process then completes its work, and returns from

main(); this usually will return into some stub code which will properly
exit the program (say, by calling theexit()system call, which traps into the OS). At this point, the OS cleans up and we are done.

The next problem with direct execution is achieving a switch between processes. Switching between processes should be simple, right? The

OS should just decide to stop one process and start another. What's the big deal? But it actually is a little bit tricky: specifically, if a process is running on the CPU, this by definition means the OS isnotrunning. If the OS is not running, how can it do anything at all? (hint: it can't) While this sounds almost philosophical, it is a real problem: there is clearly no way for the OS to take an action if it is not running on the CPU. Thus we arrive at the crux of the problem.

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Problem 2 Switching Between Processes Part 2

6.3 Problem #2: Switching Between Processes (Part 2)

THECRUX: HOWTOREGAINCONTROLOFTHECPU

How can the operating systemregain controlof the CPU so that it can switch between processes?

A Cooperative Approach: Wait For System Calls

One approach that some systems have taken in the past (for example, early versions of the Macintosh operating system [M11], or the old Xerox

Alto system [A79]) is known as thecooperativeapproach. In this style, the OStruststhe processes of the system to behave reasonably. Processes that run for too long are assumed to periodically give up the CPU so that the OS can decide to run some other task.

Thus, you might ask, how does a friendly process give up the CPU in this utopian world? Most processes, as it turns out, transfer control of the CPU to the OS quite frequently by makingsystem calls, for example, to open a file and subsequently read it, or to send a message to another machine, or to create a new process. Systems like this often include an

TIP: DEALINGWITHAPPLICATIONMISBEHAVIOR

Operating systems often have to deal with misbehaving processes, those that either through design (maliciousness) or accident (bugs) attempt to do something that they shouldn't. In modern systems, the way the OS tries to handle such malfeasance is to simply terminate the offender. One strike and you're out! Perhaps brutal, but what else should the OS do when you try to access memory illegally or execute an illegal instruction?

explicityieldsystem call, which does nothing except to transfer control to the OS so it can run other processes.

Applications also transfer control to the OS when they do something illegal. For example, if an application divides by zero, or tries to access memory that it shouldn't be able to access, it will generate atrapto the

OS. The OS will then have control of the CPU again (and likely terminate the offending process).

Thus, in a cooperative scheduling system, the OS regains control of the CPU by waiting for a system call or an illegal operation of some kind to take place. You might also be thinking: isn't this passive approach less than ideal? What happens, for example, if a process (whether malicious, or just full of bugs) ends up in an infinite loop, and never makes a system call? What can the OS do then?

A Non-Cooperative Approach: The OS Takes Control

Without some additional help from the hardware, it turns out the OS can't do much at all when a process refuses to make system calls (or mistakes) and thus return control to the OS. In fact, in the cooperative approach, your only recourse when a process gets stuck in an infinite loop is to resort to the age-old solution to all problems in computer systems:reboot the machine. Thus, we again arrive at a subproblem of our general quest to gain control of the CPU.

THECRUX: HOWTOGAINCONTROLWITHOUTCOOPERATION

How can the OS gain control of the CPU even if processes are not being cooperative? What can the OS do to ensure a rogue process does not take over the machine?

The answer turns out to be simple and was discovered by a number of people building computer systems many years ago: atimer interrupt

[M+63]. A timer device can be programmed to raise an interrupt every so many milliseconds; when the interrupt is raised, the currently running process is halted, and a pre-configuredinterrupt handlerin the OS runs.

At this point, the OS has regained control of the CPU, and thus can do what it pleases: stop the current process, and start a different one.

TIP: USETHETIMERINTERRUPTTOREGAINCONTROL

The addition of atimer interruptgives the OS the ability to run again on a CPU even if processes act in a non-cooperative fashion. Thus, this hardware feature is essential in helping the OS maintain control of the machine.

As we discussed before with system calls, the OS must inform the hardware of which code to run when the timer interrupt occurs; thus, at boot time, the OS does exactly that. Second, also during the boot sequence, the OS must start the timer, which is of course a privileged operation. Once the timer has begun, the OS can thus feel safe in that control will eventually be returned to it, and thus the OS is free to run user programs. The timer can also be turned off (also a privileged operation), something we will discuss later when we understand concurrency in more detail.

Note that the hardware has some responsibility when an interrupt oc-

curs, in particular to save enough of the state of the program that was running when the interrupt occurred such that a subsequent return-fromtrap instruction will be able to resume the running program correctly.

This set of actions is quite similar to the behavior of the hardware during an explicit system-call trap into the kernel, with various registers thus getting saved (e.g., onto a kernel stack) and thus easily restored by the

return-from-trap instruction.

Saving and Restoring Context

Now that the OS has regained control, whether cooperatively via a system call, or more forcefully via a timer interrupt, a decision has to be made: whether to continue running the currently-running process, or switch to a different one. This decision is made by a part of the operating system known as thescheduler; we will discuss scheduling policies in great detail in the next few chapters.

If the decision is made to switch, the OS then executes a low-level piece of code which we refer to as acontext switch. A context switch is conceptually simple: all the OS has to do is save a few register values for the currently-executing process (onto its kernel stack, for example) and restore a few for the soon-to-be-executing process (from its kernel stack). By doing so, the OS thus ensures that when the return-from-trap instruction is finally executed, instead of returning to the process that was running, the system resumes execution of another process.

To save the context of the currently-running process, the OS will execute some low-level assembly code to save the general purpose registers,

PC, as well as the kernel stack pointer of the currently-running process, and then restore said registers, PC, and switch to the kernel stack for the soon-to-be-executing process. By switching stacks, the kernel enters the initialize trap table remember addresses of...

syscall handler timer handler start interrupt timer start timer interrupt CPU in X ms

OS @ run Hardware Program

(kernel mode) (user mode)

Process A

...
timer interrupt save regs(A) to k-stack(A) move to kernel mode jump to trap handler

Handle the trap

Callswitch()routine save regs(A) to proc-struct(A) restore regs(B) from proc-struct(B) switch to k-stack(B)

return-from-trap (into B)
restore regs(B) from k-stack(B) move to user mode jump to B's PC

Process B

...

Figure 6.3:Limited Direct Execution Protocol (Timer Interrupt) call to the switch code in the context of one process (the one that was interrupted) and returns in the context of another (the soon-to-be-executing one). When the OS then finally executes a return-from-trap instruction, the soon-to-be-executing process becomes the currently-running process.

And thus the context switch is complete.

A timeline of the entire process is shown in Figure 6.3. In this example,

Process A is running and then is interrupted by the timer interrupt. The hardware saves its registers (onto its kernel stack) and enters the kernel (switching to kernel mode). In the timer interrupt handler, the OS decides to switch from running Process A to Process B. At that point, it calls the

switch()routine, which carefully saves current register values (into the

process structure of A), restores the registers of Process B (from its process structure entry), and thenswitches contexts, specifically by changing the stack pointer to use B's kernel stack (and not A's). Finally, the OS returnsfrom-trap, which restores B's registers and starts running it.

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Problem 2 Switching Between Processes Part 3

6.3 Problem #2: Switching Between Processes (Part 3)

Note that there are two types of register saves/restores that happen

during this protocol. The first is when the timer interrupt occurs; in this case, theuser registersof the running process are implicitly saved by the hardware, using the kernel stack of that process. The second is when the

OS decides to switch from A to B; in this case, thekernel registersare ex- 2 # 3 # Save current register context in old 4 # and then load register context from new.

5 .globl swtch 6 swtch:

7 # Save old registers 8 movl 4(%esp), %eax # put old ptr into eax 9 popl 0(%eax) # save the old IP 10 movl %esp, 4(%eax) # and stack 11 movl %ebx, 8(%eax) # and other registers 12 movl %ecx, 12(%eax) 13 movl %edx, 16(%eax) 14 movl %esi, 20(%eax) 15 movl %edi, 24(%eax) 16 movl %ebp, 28(%eax) 17 18 # Load new registers 19 movl 4(%esp), %eax # put new ptr into eax 20 movl 28(%eax), %ebp # restore other registers 21 movl 24(%eax), %edi 22 movl 20(%eax), %esi 23 movl 16(%eax), %edx 24 movl 12(%eax), %ecx 25 movl 8(%eax), %ebx 26 movl 4(%eax), %esp # stack is switched here 27 pushl 0(%eax) # return addr put in place 28 ret # finally return into new ctxt

Figure 6.4:The xv6 Context Switch Code plicitly saved by thesoftware(i.e., the OS), but this time into memory in the process structure of the process. The latter action moves the system from running as if it just trapped into the kernel from A to as if it just trapped into the kernel from B.

To give you a better sense of how such a switch is enacted, Figure 6.4 shows the context switch code for xv6. See if you can make sense of it (you'll have to know a bit of x86, as well as some xv6, to do so). The contextstructuresoldandneware found in the old and new process's process structures, respectively.

6.4 Worried About Concurrency?

Some of you, as attentive and thoughtful readers, may be now thinking: "Hmm... what happens when, during a system call, a timer interrupt occurs?" or "What happens when you're handling one interrupt and another one happens? Doesn't that get hard to handle in the kernel?" Good questions -- we really have some hope for you yet!

The answer is yes, the OS does indeed need to be concerned as to what happens if, during interrupt or trap handling, another interrupt occurs.

This, in fact, is the exact topic of the entire second piece of this book, on concurrency; we'll defer a detailed discussion until then.