`f : A → B`

to a value of its argument type A, it always produces a value of its result type B. There is no notion of definedness built into the logic, which can cause some confusion when using higher order logic formalizations of mathematical functions that are naturally partial. For example, all of the following are well-typed higher order logic terms:
`0 - 1 : natural`

`1 mod 0 : natural`

`1 / 0 : real`

`head [] : A list`

Informally, we call these *undefined terms* because they don’t have a defined mathematical value, but formally speaking they must all have a value in higher order logic. It is tempting to ask what undefined terms evaluate to, but since we are working in a theorem prover rather than a programming language there isn’t a notion of evaluation. A better question is to ask what properties are true of undefined terms, and the answer is that you can prove many theorems about them:

`0 * (0 - 1) = 0`

`(1 mod 0) - (1 mod 0) = 0`

`(0 - 1) + (1 mod 0) = (1 mod 0) + (0 - 1)`

`2 * (1 / 0) = (1 / 0) + (1 / 0)`

`head [] = head []`

Can you spot the common pattern in all the above theorems involving undefined terms? They are all true of any element of the same type as the undefined term:

`∀n : natural. 0 * n = 0`

`∀n : natural. n - n = 0`

`∀m n : natural. m + n = n + m`

`∀x : real. 2 * x = x + x`

`∀x : A. x = x`

In other words, proving theorems like this about undefined terms gives us no more information than is already provided by their type. However, depending on how we have formalized the partial mathematical function we may be able to also prove *accidental theorems* about undefined terms that give us more information about them. For example, consider formalizing natural number subtraction with the following defining equations:

`∀m. m - 0 = m`

`∀m n. suc m - suc n = m - n`

This seems harmless enough: we cannot use these equations to prove that an undefined subtraction term such as `0 - 1`

is equal to a particular natural number. However, we can instantiate the second equation to prove that two undefined subtraction terms are equal:

`1 - 2 = 0 - 1`

This is an accidental theorem, which we can check by replacing the undefined terms with arbitrary natural numbers:

`∀m n : natural. m = n`

Since this generalized theorem is clearly false, it demonstrates that the accidental theorem reveals more information about the undefined subtraction terms than is already known by their type.

Fortunately, it is usually straightforward to eliminate accidental theorems. In the case of subtraction we simply replace the second defining equation with the following:

`∀m n. n ≤ m ⇒ suc m - suc n = m - n`

After this change there can be no more accidental theorems about undefined subtraction terms, and we can derive no information about them other than they have natural number type.

Suppose we have carefully formalized our partial mathematical functions to avoid any accidental theorems. Then we expect every theorem about the `mod`

function to have a side-condition that ensures the second argument is non-zero, such as the following:

`∀n a b. ¬(n = 0) ⇒ (a * b) mod n = ((a mod n) * (b mod n)) mod n`

`∀n a. a < n ⇒ a mod n = a`

The reason for this side condition is that by design we have no information about the value of undefined `mod`

terms, and so to prove properties of `mod`

we have to rule out the case where the second argument is zero. This is satisfying, because the mathematical version of these theorems have the same side-condition to ensure that all terms are defined.

However, there are some *surprising theorems* in the formalization that do not have this side-condition, such as the following:

`∀n a b. (a * n + b) mod n = b mod n`

Why doesn’t this theorem need a side-condition that n is non-zero to avoid undefined `mod`

terms? Well, if n is zero, then

`a * n + b = a * 0 + b = 0 + b = b`

reducing the left hand side of the surprising theorem to be exactly the same as the right hand side:

`b mod 0 = b mod 0`

Since this is true of every natural number, we can prove it for these undefined `mod`

terms even though we have no information about their value.

The existence of surprising theorems is a consequence of formalizing partial mathematical functions in higher order logic, in which there is no notion of definedness and every function is total. Even taking care to avoid accidental theorems, this fundamental logical mismatch leads to differences between mathematical theorems and their formalized versions.

In most situations this does not cause any problems, and it is merely a curiosity that certain theorems do not have the expected side-conditions. But in some circumstances, such as the need to export formalized proofs to environments that do have a notion of definedness, it may be critical to maintain a close correspondence between the mathematical theorems and their formalized versions. Happily, since the surprising feature of surprising theorems is that they are stronger than expected, we can simply choose to formalize deliberately weaker versions with logically unnecessary side-conditions, such as the following:

`∀n a b. ¬(n = 0) ⇒ (a * n + b) mod n = b mod n`

In this way we can maintain a close correspondence between the syntax of mathematical theorems and their formalized versions, if that be necessary.

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**Prototyping:** I begin quite normally by writing some code in Haskell, testing out different ways of solving some problem that I’m interested in. In this step I’m mainly focused on achieving a clean design with reasonable performance, and wherever possible I use existing Haskell packages (but restricted to those that have been developed with the same software development methodology). I also write QuickCheck tests to check various properties of the functions in the design. At this point I’m essentially done with the traditional coding phase, though usually I spend what might be perceived as an excessive amount of time polishing to make the design and properties as simple and clear as possible.

**Verification:** Next I use the HOL Light interactive theorem prover to develop a logical theory that (i) defines the types and functions in my Haskell design, and (ii) proves the QuickCheck properties as theorems that follow from the definitions. If the Haskell package depends on other packages that follow the same software development methodology then their contents must also be formalized in theories, and the logical definitions can build on each other in the same way as the Haskell definitions. This step is the most labour-intensive, and often some degree of creativity is required to formulate properties of recursive functions that can be proved by induction. This is where the effort spent polishing the prototype repays itself, since any unnecessary complexity in the design will inevitably thread its way through the proof of many properties.

**Export:** The final step is to use the opentheory tool to export the logical theory as a Haskell package. I first package up the HOL Light theory in OpenTheory format, by rendering its proof in a low-level article format and adding some meta-data such as the package name and the list of OpenTheory packages it depends on. There is also some Haskell-specific meta-data, such as the mapping from OpenTheory symbol names to Haskell names, and selecting which theorems should be exported as QuickCheck properties. Once that is complete the opentheory tool can automatically export the OpenTheory package as a Haskell package (in a form ready to be uploaded to Hackage).

At this point you might reasonably ask: why develop a Haskell package, throw it away, and then develop a logical theory only to export the same Haskell package again? Excellent question, and the rest of this post is devoted to answering it, but let’s begin with a depressing truth: it’s not the proof. Despite the countless hours and myriad creative insights you poured into that final Q.E.D., no one wants to see or even hear about your proof. Formal verification proofs are write-only. Sorry about that.

**Specification:** The proofs might be irrelevant, but the proved theorems constitute a precise specification of the Haskell package, grounded all the way down to standard mathematical objects formalized in higher order logic. I view Haskell as a computational platform for higher order logic (rather than higher order logic as a verification platform for Haskell), a perspective that emphasizes the consistent set of logical theories underlying the Haskell package that defines its semantics. A practical consequence of this is that any differences in the behaviour of the Haskell package and its higher order logic semantics (for example, due to bugs in the export code) must be fixed by changing the Haskell computational platform to conform to its higher order logic semantics.

Let’s take a look at an example specification for an extended GCD function `egcd`

defined in the Haskell package opentheory-divides:

`egcd :: Natural -> Natural -> (Natural, (Natural, Natural))`

The following theorem about `egcd`

is one of many proved in the corresponding logical theory natural-divides:

`∀a b g s t. ¬(a = 0) ∧ egcd a b = (g,(s,t)) ⇒ t * b + g = s * a`

This says that the three components of the `egcd`

result (g, s and t) will satisfy a particular equation with its arguments (a and b), *so long as the first argument a is non-zero*. Side-conditions like this don’t fit naturally into the Haskell type system, and so the burden falls on programmers to discover them and code around them when necessary. This may be one reason why programmers prefer to write their own software components (with side-conditions tailored to the application) rather than reuse existing components (with unknown and/or inconvenient side-conditions). Using this software development methodology I shift the burden of discovering side-conditions to HOL Light, which will force me to prove that the first argument of `egcd`

is non-zero in any code that relies on the above property of the result. If the code just uses the property that the result g of `egcd`

is the greatest common divisor then HOL Light will understand that no side-condition proof is necessary, because another theorem in the logical theory proves that this is true for all arguments to `egcd`

:

`∀a b. fst (egcd a b) = gcd a b`

**Proof:** It may be surprising to learn that verified Haskell packages can be useful as a building block to formalize mathematics in higher order logic. Once we have defined Haskell types and functions in the logic and proved that they satisfy their specification, we can build on this formal development to prove mathematical theorems that do not refer to any Haskell symbols. To continue the `egcd`

example, suppose we have two natural numbers a and b (b is greater than 1) that are coprime (i.e., `gcd a b = 1`

). Then we can reduce modulo b both sides of the first property of `egcd`

to show the existence of a multiplicative inverse for a modulo b:

`∀a b. 1 < b ∧ gcd a b = 1 ⇒ ∃s. (s * a) `mod` b = 1`

This is a non-trivial higher order logic theorem which refers only to standard mathematical concepts, but its proof rests on the properties of the verified Haskell `egcd`

function.

**Maintenance:** Software maintenance is a pain, and it is here that this software development methodology really shines (and who knows, maybe eventually cause the tortoise to pass the hare?). Formally verified software offers a new capability: by examining the logical theory of a Haskell package, we can see exactly which properties of dependent packages it relies on. We can use this information to automatically check whether or not a new version of a dependent package will preserve functional correctness: if so then extend the acceptable version range; and if not then report an error to the maintainer. The opentheory tool uses this technique to automatically generate maximal acceptable version ranges for dependent packages whenever it exports Haskell packages, removing another cognitive burden from the programmer. Earlier I said that formal verification proofs were write-only, but really all I meant is that humans won’t read them: as this version analysis shows machines are quite happy to read them and extract useful information.

**Fun:** Actually I retract the whole depressing truth: it is totally about the proof. Not for showing to people (humans still won’t read formal verification proofs) but rather for the pure fun of writing them. Those myriad creative insights you made over the course of countless hours? For me formal verification is as immersive as a video game, and my high scores come with a new artefact to play with and build upon. Free toy inside!

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`gcd 15 6 == 3`

gcd 21 10 == 1

gcd 10 0 == 10

gcd 0 0 == 0 -- by convention

Two numbers that have a gcd of 1 (like 21 and 10) are said to be *coprime*. Although in general it’s a hard problem to factor an integer into divisors

`91 == 7 * 13`

69 == 3 * 13

it’s an easy problem to compute the gcd of two integers

`gcd 91 69 == 13`

and if they are not coprime then a non-trivial divisor of both is revealed.

Here is a simple recursive definition that efficiently computes gcd:

`gcd :: Integer -> Integer -> Integer`

gcd a 0 = a

gcd a b = gcd b (a `mod` b)

How would we check that this definition always computes the gcd of two non-negative integers? If we define a divides relation

`divides :: Integer -> Integer -> Bool`

divides 0 b = b == 0

divides a b = b `mod` a == 0

then we can easily test that gcd returns *some* common divisor:

`commonDivisorTest :: Integer -> Integer -> Bool`

commonDivisorTest a b = let g = gcd a b in divides g a && divides g b

But how to check that gcd returns the *greatest* common divisor? Consider the linear combination g of a and b

`g = s * a + t * b`

where the coefficients s and t are arbitrary integers. Any common divisor of a and b must also be a divisor of g, and so if g is itself a common divisor of a and b then it must be the greatest common divisor. The *extended* gcd algorithm computes the coefficients s and t together with the gcd:

`egcd :: Integer -> Integer -> (Integer,(Integer,Integer))`

egcd a 0 = (a,(1,0))

egcd a b =

(g, (t, s - (a `div` b) * t))

where

(g,(s,t)) = egcd b (a `mod` b)

which allows us to test that egcd returns the greatest common divisor:

`greatestCommonDivisorTest :: Integer -> Integer -> Bool`

greatestCommonDivisorTest a b =

let (g,(s,t)) = egcd a b in

divides g a && divides g b && s * a + t * b == g

The coefficients s and t are useful for much more than checking the gcd result: for any coprime a and b we have

`s * a + t * b == 1`

(s * a) `mod` b == 1 -- assuming b > 1

(t * b) `mod` a == 1 -- assuming a > 1

which shows that s is the multiplicative inverse of a modulo b, and t is the multiplicative inverse of b modulo a.

The extended gcd algorithm is a classic integer algorithm, but what if we wanted a version that computed over the type of natural numbers? This would be useful in environments in which unsigned arithmetic is more natural and/or efficient than signed arithmetic (one such environment being the natural number theory of an interactive theorem prover, which is where this need arose).

The immediate problem is that one of the s and t coefficients returned by the extended gcd algorithm is usually a negative number, which is to be expected since they must satisfy

`s * a + t * b == g`

where g is generally less than both a and b. The solution is to modify the specification and ask for natural number coefficients s and t that satisfy

`s * a == t * b + g`

Fortunately for us, it turns out that such coefficients always exist except when a is zero (and b is non-zero).

Unfortunately for us, our modified specification is no longer symmetric in a and b. Given coprime a and b as input, s is the multiplicative inverse of a modulo b as before, but t is now the *negation* of the multiplicative inverse of b modulo a (the additive inverse of the multiplicative inverse—funny). This asymmetry breaks the recursion step of the extended gcd algorithm, because there is no longer a way to derive coefficients s and t satisfying

`s * a == t * b + g`

from coefficients s’ and t’ satisfying

`s' * b == t' * (a `mod` b) + g`

[If you try you’ll discover that a and b are the wrong way around.] This problem is solved by unwinding the recursion so that the algorithm makes two reductions before calling itself recursively, giving coefficients s’ and t’ satisfying

`s' * (a `mod` b) == t' * (b `mod` (a `mod` b)) + g`

Although this is expression is more complicated, it is now possible to define coefficients

`s = s' + (b `div` (a `mod` b)) * t'`

t = t' + (a `div` b) * s

which satisfy the desired equation

`s * a == t * b + g`

Here is the finished implementation:

`egcd :: Natural -> Natural -> (Natural,(Natural,Natural))`

egcd a 0 = (a,(1,0))

egcd a b =

if c == 0

then (b, (1, a `div` b - 1))

else (g, (u, t + (a `div` b) * u))

where

c = a `mod` b

(g,(s,t)) = egcd c (b `mod` c)

u = s + (b `div` c) * t

This satisfies a modified test confirming that it produces a greatest common divisor (for non-zero a):

`greatestCommonDivisorTest :: Natural -> Natural -> Bool`

greatestCommonDivisorTest a b =

let (g,(s,t)) = egcd a b in

divides g a && divides g b && s * a == t * b + g

And we can also check the upper bounds on the s and t coefficients (for non-zero a and b at least 2):

`coefficientBoundTest :: Natural -> Natural -> Bool`

coefficientBoundTest a b =

let (_,(s,t)) = egcd a b in

s < b && t < a

The *Chinese Remainder Theorem* states that for coprime a and b, given an x and y satisfying

`x < a`

y < b

there is a unique n satisfying

`n < a * b`

n `mod` a == x

n `mod` b == y

But how to compute this n given natural number inputs a, b, x and y? We build on the natural number extended gcd algorithm, of course:

`chineseRemainder :: Natural -> Natural -> Natural -> Natural -> Natural`

chineseRemainder a b =

\x y -> (x * tb + y * sa) `mod` ab

where

(_,(s,t)) = egcd a b

ab = a * b

sa = s * a

tb = (a - t) * b

The λ expression in the body maximizes the precomputation on a and b, allowing efficient Chinese remaindering of different x and y inputs once the a and b inputs have been fixed. As with all computation over the natural numbers we must be careful with the subtraction operation `(a - t)`

to ensure that the first argument is always larger than the second, but here we know from the `coefficientBoundTest`

that t is always less than a.

All of the natural number computations in this blog post have been formally verified and are available as the Haskell package opentheory-divides.

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**Creative**: Or it should be automated. Sometimes automating non-creative work can itself be creative work.**Useful**: Or it will never see the light of day, because it doesn’t solve a problem that is causing someone pain.**Pure**: Or it will accumulate fat and inevitably collapse under its own weight.

Work with these three attributes I will seek out, I will give the time and attention to develop, I will ship the product **and** maintain it.

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Well, it’s one thing to look for knight’s tours by hand, but quite another to search for them with a computer. It is actually a fun programming exercise to search for knight’s tours: it is possible to use a standard backtracking algorithm, but it needs a couple of tricks to make it work. Firstly, if you have a choice of squares to jump to, then you should consider them in ascending order of *accessibility* (i.e., from how many unvisited squares could you jump to this square). This helps to ensure that you don’t waste last chances to visit squares. However, even with this trick, it is possible to make mistakes early on in the search that you can never recover from, and end up lost forever in a maze with 10^{51} turns and no exits. An ugly way to fix this is to assign a move budget—I chose 100 times the number of squares—and if you don’t find a solution within this many jumps then give up and start again from an empty board.

Armed with a fast computer, a search algorithm and a couple of tricks, I was finally able to realize a childhood dream and find a knight’s tour:

Note that we always begin our tours in the bottom right corner, so that the search program only has to work out how to get into and out of 3 corners: that time spent filling squared notebooks taught me something!

Once the initial problem is solved, I naturally start wondering whether we can find knight’s tours on larger boards. It is easy to modify the search program to change the board dimensions from an 8×8 square to an *m*×*n* rectangle, but as the board size increases the time it takes to find a knight’s tour grows rapidly:

Board size (n×n) |
8 |
10 |
12 |
14 |
16 |
18 |
20 |
22 |
24 |
26 |
28 |
30 |

Search time (ms) |
1 | 3 | 7 | 11 | 26 | 51 | 102 | 225 | 401 | 701 | 1,074 | 1,738 |

Extrapolating these results to a 100×100 board, it seems that my search program would take over a thousand years to find a knight’s tour! Clearly I need a new approach to find knight’s tours for large boards.

Our 8×8 knight’s tour above looks like a little circuit, and drawn the same way a 100×100 knight’s tour would look like a huge circuit. But let’s take some inspiration from the way that huge circuits are designed in practice: not all at once, but rather by wiring together smaller circuits. Consider this special 6×6 knight’s tour:

One reason this knight’s tour is special is because it has no start and end: it is a loop! There are probably clever ways to find knight’s tours that are loops, but I took a simple-minded approach. I modified my program to keep searching for knight’s tours until it found one where the end was a knight’s jump away from the start, and then I connected the ends of the tour to make a loop.

There is another reason that this knight’s tour is special, which is not so easy to see. You can put two of them together:

then switch two of the jumps around to connect the tours together:

and the result is a knight’s tour on a 12×6 board. To see that the result is indeed a tour, let’s visualize the two separate knight’s tours as loops using a schematic diagram:

And now when we perform the connection step the schematic diagram makes it clear that we have created one big loop from the two separate loops:

The same thing also works by vertically stacking the special 6×6 knight’s tour:

and connecting them like so:

I’ve shown how to connect special 6×6 knight’s tours together, and the next step is to connect them in an *m*×*n* grid to make a knight’s tour of a 6*m*×6*n* board. Connecting them together in a grid requires a little bit of care to avoid fragmenting the knight’s tour into disconnected paths. Here’s a schematic diagram for a 2×2 grid illustrating how tours can be fragmented into disconnected paths:

One way to avoid tour fragmentation is to connect along some path through the grid that visits every 6×6 knight’s tour. I chose a spiral, starting at the bottom right of the grid, travelling anti-clockwise and ending in the middle. Here is the schematic diagram for a spiral path on a 3×3 grid:

And here’s the resulting knight’s tour on an 18×18 board:

At this stage we have the machinery to construct knight’s tours for arbitrarily large boards of dimension 6*m*×6*n*. But what about boards of other dimensions, such as the 100×100 board that we set as our target? These boards can be handled by searching for more special tours of dimensions *m*×*n*, where 6 ≤ *m,n* ≤ 11, and using these in place of 6×6 knight’s tours on the bottom and right edges of the grid. For example, here is a knight’s tour of a 23×23 board:

This has the same 3×3 grid pattern as the 18×18 knight’s tour, but with larger pieces around the right and bottom edges. Also, the knight’s tour in the bottom right corner is not a loop—you can see one of its ends in the bottom right corner—so the schematic diagram for the assembled tour ends up looking a little different:

In fact, it is impossible to find a knight’s tour that is a loop on this board, because it has an odd number of squares. To see why, imagine the squares of the board painted light and dark colors in a checkerboard pattern:

As a knight jumps around the board it alternates between light and dark squares, and a tour of a board with an odd number of squares will contain an even number of jumps. This means the start and end squares of the tour will have the same color, and so there’s no way to jump from the end square of the tour back to the start square to close the loop.

Using these techniques you can easily assemble knight’s tours for large boards: you can view a tour of a 100×100 board or even assemble your own knight’s tour.

But these techniques only work for boards where the width and height are at least 6: there are also an infinite number of *thin boards* that might contain knight’s tours:

Can you think of techniques for assembling knight’s tours on thin boards? ♘

The performance results were gathered using a MacBook Air with a 1.8GHz Intel Core i7 processor and 4Gb of RAM, running OS X 10.7.5.

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To illustrate the explicit laziness technique, we’ll use it to optimize an Haskell program for computing prime numbers based on a sieve method. This is the central `Sieve`

data structure:

`newtype Sieve =`

Sieve { unSieve :: (Natural,[(Natural,Natural)]) }

Values of type `Natural`

are non-negative integers of unbounded size. Every value `Sieve (n,ps)`

of type `Sieve`

must satisfy the invariant that `map fst ps`

is the list of all prime numbers up to the *perimeter* `n`

in ascending order, and the *counter* `k`

in every element `(p,k)`

of the list `ps`

satisfies `k == n `mod` p`

. The following examples are valid values of type `Sieve`

:

`Sieve (9,[(2,1),(3,0),(5,4),(7,2)])`

`Sieve (10,[(2,0),(3,1),(5,0),(7,3)])`

`Sieve (11,[(2,1),(3,2),(5,1),(7,4),(11,0)])`

The `Sieve`

API supports two operations for creating new values of type Sieve:

`initial :: Sieve`

increment :: Sieve -> (Bool,Sieve)

The `initial`

value of a `Sieve`

is simply:

`initial :: Sieve`

initial = Sieve (1,[])

The `increment`

operation takes a value `Sieve (n,_)`

and returns a new value `Sieve (n + 1, _)`

, together with a boolean indicating whether the new perimeter `n + 1`

is a prime number:

`increment :: Sieve -> (Bool,Sieve)`

increment =

\s ->

let (n,ps) = unSieve s in

let n' = n + 1 in

let ps' = inc ps in

let b = check ps' in

let ps'' = if b then ps' ++ [(n',0)] else ps' in

(b, Sieve (n', ps''))

where

inc = map (\(p,k) -> (p, (k + 1) `mod` p))

check = all (\(_,k) -> k /= 0)

The `increment`

operation works by adding 1 to the perimeter `n <- n + 1`

and every counter `k <- (k + 1) `mod` p`

. It then checks whether any counter is zero, which happens when the perimeter is divisible by a known prime. If no counter is zero then the perimeter is a new prime, and so `(n,0)`

is added to the end of the list.

How can we optimize this program? Well, there are certainly many ways that we can help the compiler produce efficient target code, but before we get into that we should consider algorithmic improvements. For example, are we performing any unnecessary computation? It seems that we are, because we increment every counter even though the boolean result only needs to know whether any counter is zero. We could stop at the first zero counter and just make a note (called a *closure*) that the other counters need to be incremented later. This is an instance of lazy evaluation, and since Haskell is a lazy language this optimization will be performed automatically.

Unfortunately, there is still a performance problem here. Consider a time in the future when we are incrementing counters and we come to a counter with an increment closure. At this point Haskell’s lazy evaluation will execute the closure to increment the counter, and then increment it for a second time to see whether it is zero. This is functionally correct, but it would be much more efficient to add 2 to the counter rather than add 1 and then add 1 again. Even worse, it turns out that the increment closures stack up on counters during evaluation, to the point that each time we discover a new prime `p`

we have to increment the counter `k`

in the final list element `(q,k)`

a total of `p - q`

times before we can check that it is non-zero and record the new prime.

How can we avoid this inefficient behavior of adding `m`

to a counter by adding 1 to it `m`

times? This is a case where Haskell’s automatic laziness is not flexible enough, because whenever evaluation encounters a value with a closure, the only possible course of action is to execute the closure and then continue evaluation. There is no way to examine the closure to find out what computation it will perform. However, we can manually simulate lazy evaluation by storing data structures together with explicit notes saying what computation needs to be done before they can be used.

Here is how explicit laziness works in our example. First, we extend the Sieve data structure as follows:

`newtype Sieve =`

Sieve { unSieve :: (Natural,[(Natural,(Natural,Natural))]) }

This is exactly the same as before, except that each list element `(p,(k,j))`

has an extra natural number component `j`

, which is a note that `j`

must be added to every counter following this one in the list.

The definition of the `initial`

value remains the same:

`initial :: Sieve`

initial = Sieve (1,[])

However, the `increment`

operation now takes into account the explicit `j`

notes:

`increment :: Sieve -> (Bool,Sieve)`

increment =

\s ->

let (n,ps) = unSieve s in

let n' = n + 1 in

let (b,ps') = inc n' 1 ps in

(b, Sieve (n',ps'))

where

inc n _ [] = (True, (n,(0,0)) : [])

inc n i ((p,(k,j)) : ps) =

let k' = (k + i) `mod` p in

let j' = j + i in

if k' == 0 then (False, (p,(0,j')) : ps)

else let (b,ps') = inc n j' ps in (b, (p,(k',0)) : ps')

The `inc`

and `check`

functions in the previous implementation have been combined into one `inc`

function which takes a parameter `i`

recording how much to increment counters. Initially `i`

is set to 1, but it absorbs the `j`

values for the counters it visits. As soon a zero counter is discovered, the i parameter is added to the current j value **and the rest of the list is untouched**. This last point is crucial: explicit laziness eliminates the overhead of closures that are normally generated to update data structures not directly required for the result. For comparison with the version of the program without explicit laziness, here are the `Sieve`

values with perimeters 9, 10 and 11:

`Sieve (9,[(2,(1,0)),(3,(0,2)),(5,(2,0)),(7,(0,0))])`

`Sieve (10,[(2,(0,1)),(3,(0,2)),(5,(2,0)),(7,(0,0))])`

`Sieve (11,[(2,(1,0)),(3,(2,0)),(5,(1,0)),(7,(4,0)),(11,(0,0))])`

Values of the new version of the `Sieve`

type are much harder to read than values of the original version, because the `j`

values have to be added to every counter following it in the list. This complicates the invariant for `Sieve`

values, but is actually a measure of the efficiency gained. For example, the `j`

value of 2 in the `Sieve`

value with perimeter 9 line replaces 4 counter increments in the original program: 2 for the 5 counter and another 2 for the 7 counter.

Even though explicit laziness is an algorithmic optimization, and there remain many optimizations that can be applied to both versions of the example sieve program, it is nevertheless interesting to gauge the performance effects of explicit laziness using a standard compiler. The following table presents the time taken to compute the *N*th prime using both versions of the example sieve program, for various values of *N*:

Nth prime |
100 |
200 |
400 |
800 |
1600 |
3200 |

Without explicit laziness |
0.007s | 0.019s | 0.068s | 0.284s | 1.169s | 5.008s |

With explicit laziness |
0.004s | 0.008s | 0.025s | 0.083s | 0.345s | 1.495s |

It is gratifying to see that the version of the example sieve program with explicit laziness is about 3x faster on this small experiment.

In this post we have described explicit laziness, a useful technique for optimizing functional programs. Like all optimization techniques its effectiveness depends on the context: automatic laziness is more flexible and guarantees that closures will be evaluated at most once, so should be the default. However, the judicious introduction of explicit laziness can eliminate closures and allow future computations to be combined into more efficient variants before executing them (a limited form of computation on computations).

The code for the optimized version of the example sieve program is available at Hackage: this verified Haskell package is automatically generated from an OpenTheory package. The performance results were gathered using the GHC 7.4.1 compiler on a MacBook Air with a 1.8GHz Intel Core i7 processor and 4Gb of RAM, running OS X 10.7.5.

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