One of the most exciting papers I’ve read in a long time is James and Sabry’s Information Effects. It starts with the hook “computation is a physical process which, like all other physical processes, is fundamentally reversible,” and it goes from there. If that doesn’t immediately pull you in, perhaps some of the subsequent PL jargon will – it promises a “typed, universal, and reversible computation model in which information is treated as a linear resource”.

I don’t know about you, but I was positively shaking with anticipation at this point. That’s one heck of an abstract.

After some philosophy and overview of the paper, James and Sabry dive into the appetizer in a section titled “Thermodynamics of Computation and Information”. They give the following definition:

DEFINITION 2.2 (Entropy of a variable). Let \(b\) be a (not necessarily finite) type whose values are labeled \(b_1\), \(b_2\), \(\ldots\). Let \(\xi\) be a random variable of type \(b\) that is equal to \(b_i\) with probability \(p_i\). The entropy of \(\xi\) is defined as \(- \sum p_i \log{p_i}\).

and the following, arguably less inspired definition:

DEFINITION 2.3 (Output entropy of a function). Consider a function f : a -> b where b is a (not necessarily finite) type whose values are labeled \(b_1\), \(b_2\), \(\ldots\). The output entropy of the function is given by \(- \sum q_j \log{q_j}\) where \(q_j\) indicates the probability of the function to have value \(b_j\).

We can say now that a function is reversible if and only if the entropy of its arguments is equal to the entropy of its output. Which is to say that the gain in entropy across the function is 0.

Of course, as astute students of mathematics we know that reversibility of a function is equivalent to whether that function is an isomorphism. While this is how we will prefer to think of reversibility, the definition in terms of entropy brings up interesting questions of pragmatics that we will get to later.

James et al. present the following language, which we have reproduced here translated into Haskell. The language is first order, and so we will ignore function types, giving us the types:

{-# LANGUAGE TypeOperators #-} {-# LANGUAGE NoImplicitPrelude #-} -- Corresponds to Haskell '()' type data U = U -- Corresponds to Haskell 'Either' type data a + b = InL a | InR a -- Corresponds to Haskell '(,)' type data a * b = Pair a b a b

The language presented is based around the notion of type isomorphisms, and so in order to model this language in Haskell, we’ll need the following types:

data a <=> b = Iso { run :: a -> b , rev :: b -> a }

This type a <=> b represents an isomorphism between type a and type b , as witnessed by a pair of functions to and from . This probably isn’t the best encoding of an isomorphism, but for our purposes it will be sufficient.

James and Sabry present the following axioms of their language:

swapP :: a + b <=> b + a assocP :: a + (b + c) <=> (a + b) + c (bc)(ab) unite :: U * a <=> a swapT :: a * b <=> b * a assocT :: a * (b * c) <=> (a * b) * c (bc)(ab) distrib :: (a + b) * c <=> (a * c) + (b * c) (ab)(ac)(bc) id :: a <=> a

The implementations of these terms are all trivial, being that they are purely syntactic isomorphisms. They will not be reproduced here, but can be found in the code accompanying this post. The motivated reader is encouraged to implement these for themself.

With the terms of our algebra out of the way, we’re now ready for the operators. We are presented with the following:

-- Isomorphisms are symmetric. sym :: (a <=> b) -> (b <=> a) (ab)(ba) -- Isomorphisms are transitive. (>>) :: (a <=> b) -> (b <=> c) -> (a <=> c) (ab)(bc)(ac) -- Products and coproducts are bifunctors. (.+) :: (a <=> c) -> (b <=> d) -> (a + b <=> c + d) (ac)(bd)(ad) (.*) :: (a <=> c) -> (b <=> d) -> (a * b <=> c * d) (ac)(bd)(ad)

It turns out that the resulting language is already surprisingly expressive. We can encode booleans:

type Bool = U + U false :: Bool true, = InL U true = InR U false

With these out of the way, James et al. show us a “one-armed if-expression”: if the incoming Bool is true , transform the a by the provided combinator:

ifthen :: (a <=> a) -> ( Bool * a <=> Bool * a) (aa)a) = distrib >> ( id .* c) .+ id >> sym distrib ifthen cdistribc)sym distrib

For syntactic convenience, we will enable rebindable syntax, allowing us to represent these chains of isomorphism transitivities with do notation. We can thus express the above more clearly as:

{-# LANGUAGE RebindableSyntax #-} ifthen :: (a <=> a) -> ( Bool * a <=> Bool * a) (aa)a) = do ifthen c distrib ( id .* c) .+ id c) sym distrib

The mystery behind naming our transitivity operator (>>) is thus explained.

But how does our ifthen combinator actually work? Recall that Bool = U + U , meaning that we can distribute the U s across the pair, giving us the type (U * a) + (U * a) . The left branch (of type U * a ) of this coproduct has an inhabitant if the incoming boolean was true .

We can thus bimap over the coproduct. Since the left case corresponds to an incoming true , we can apply an isomorphism over only that branch. Because we want to transform the incoming a by the combinator c , we then bimap over our U * a with id .* c – not touching the U but using our combinator.

Finally, we need to repackage our (U * a) + (U * a) into the correct return type Bool * a , which we can do by factoring out the a . Factoring is the inverse of distrib uting, and so we can use the sym operator to “undo” the distrib .

It’s crazy, but it actually works! We can run these things to convince ourselves. Given:

not :: Bool <=> Bool not = swapP -- move a left ('true') to a right ('false'), and vice versa. swapP

We get:

> run (ifthen not) $ Pair true false Pair true true > run (ifthen not) $ Pair false false Pair false false

Neat, no? For fun, we can also run these things backwards:

> rev (ifthen not) $ Pair true true Pair true false > rev (ifthen not) $ Pair false false Pair false false

James et al. are eager to point out that ifthen (ifthen not) :: Bool * (Bool * Bool) <=> Bool * (Bool * Bool) is the Toffoli gate – a known universal reversible gate. Because we can implement Toffoli (and due to its universalness), we can thus implement any boolean expression.

Recursion and Natural Numbers

Given two more primitives, James and Sabry show us how we can extend this language to be “expressive enough to write arbitrary looping programs, including non-terminating ones.”

We’ll need to define a term-level recursion axiom:

trace :: (a + b <=> a + c) -> (b <=> c) (ac)(bc)

The semantics of trace are as follows: given an incoming b (or, symmetrically, a c ), lift it into InR b :: a + b , and then run the given iso over it looping until the result is an InR c , which can then be returned.

Notice here that we have introduced potentially non-terminating looping. Combined with our universal boolean expressiveness, this language is now Turing-complete, meaning it is capable of computing anything computable. Furthermore, by construction, we also have the capability to compute backwards – given an output, we can see what the original input was.

You might be concerned that the potential for partiality given by the trace operator breaks the bijection upon which all of our reversibility has been based. This, we are assured is not a problem, because a divergence is never actually observed, and as such, does not technically violate the bijectiveness. It’s fine, you guys. Don’t worry.

There is one final addition we need, which is the ability to represent inductive types:

data Fix f = Fix { unFix :: f ( Fix f) } f (f) } fold :: f ( Fix f) <=> Fix f f (f) = Iso Fix unFix foldunFix

Given these things, we can define the natural numbers a little circuitously. We can define their type as follows:

type Nat = Fix (( + ) U ) ((

Constructing such things is a little tricky, however. First we’ll need a way to introduce a coproduct. The type and name of this isomorphism should be suggestive:

just :: a <=> U + a = trace $ do justtrace sym assocP >> swapP) .+ id (sym foldswapP)

just is a tricky little beast; it works by using trace to eliminate the Nat + U of a (Nat + U) + (U + a) . We can follow the derivation a little more closely:

body :: ( Nat + U ) + a <=> ( Nat + U ) + ( U + a) a) = do body -- Nat + (U + a) sym assocP .+ id -- (U + Nat) + (U + a) sym fold .+ id -- (Nat + U) + (U + a) swapP body :: a <=> U + a trace

I wish I had come up with this, because it’s quite clever. Notice however that this is a partial isomorphism; when run backwards, it will diverge in the case of InR U :: U + a .

Given just , we can now define succ :

succ :: Nat <=> Nat succ = do -- U + Nat just -- Nat fold

James et al. provide a little more machinery in order to get to the introduction of a 0:

injectR :: a <=> a + a = do injectR -- U * a sym unite .* id -- (U + U) * a just -- (U * a) + (U * a) distrib .+ unite -- a + a uniteunite

and finally:

zero :: U <=> Nat = trace $ do zerotrace id -- Nat + U -- U + Nat swapP -- Nat fold -- Nat + Nat injectR

What’s interesting here is that the introduction of 0 is an isomorphism between U and Nat , as we should expect since 0 is a constant.

Induction on Nats

The paper teases an implementation of isEven for natural numbers – from the text:

For example, it is possible to write a function even? :: Nat * Bool <=> Nat * Bool which, given inputs (n, b) , reveals whether n is even or odd by iterating not n -times starting with b . The iteration is realized using trace as shown in the diagram below (where we have omitted the boxes for fold and unfold ).

Emphasis mine. The omitted fold and unfold bits of the diagram are the actual workhorses of the isomorphism, and their omission caused me a few days of work to rediscover. I have presented the working example here to save you, gentle reader, from the same frustration.

The insight is this – our desired isomorphism has type Nat * a <=> Nat * a . Due to its universally qualified nature, we are unable to pack any information into the a , and thus to be reversible, the Nat must be the same on both sides. Since we are unable to clone arbitrary values given our axioms (seriously! try it!), our only solution is to build a resulting Nat up from 0 as we tear apart the one we were given.

We can view the a in trace :: (a + b <=> a + c) -> (b <=> c) as “scratch space” or “intermediate state”. It is clear that in order to execute upon our earlier insight, we will need three separate pieces of state: the Nat we’re tearing down, the Nat we’re building up, and the a along for the ride.

For reasons I don’t deeply understand, other than it happened to make the derivation work, we also need to introduce a unit to the input of our traced combinator.

With this line of reasoning, we have the following:

iterNat :: (a <=> a) -> ( Nat * a <=> Nat * a) (aa)a) = do iterNat step sym unite $ do trace id -- (Nat' * (Nat * a)) + (U * (Nat * a)) unite

For clarity, we’ll annotate the natural number under construction as Nat' .

When the iteration begins, our combinator receives an InR whose contents are of type U * (Nat * a) corresponding to the fact that there is not yet any internal state. From there we can factor our the Nat * a :

... $ do trace id -- (Nat' * (Nat * a)) + (U * (Nat * a)) -- (Nat' + U) * (Nat * a) sym distrib ...

All of a sudden this looks like a more tenable problem. We now have a product of (conceptually) a Maybe Nat' , the Nat being torn down, and our a . We can fold :: U + Nat <=> Nat our Nat' , which will give us 0 in the case that the state hasn’t yet been created, or \(n+1\) in the case it has.

... $ do trace id -- (Nat' * (Nat * a)) + (U * (Nat * a)) -- (Nat' + U) * (Nat * a) sym distrib >> fold) .* id -- Nat' * (Nat * a) (swapPfold) ...

The only thing left is to destruct the incoming Nat and apply our step isomorphism. We introduce a lemma to help:

swapBacT :: a * (b * c) <=> b * (a * c) (bc)(ac) = do swapBacT assocT .* id swapT sym assocT

which we can then use to move the pieces of our state and destruct the correct number:

... $ do trace id -- (Nat' * (Nat * a)) + (U * (Nat * a)) -- (Nat' + U) * (Nat * a) sym distrib >> fold) .* id -- Nat' * (Nat * a) (swapPfold) -- Nat * (Nat' * a) swapBacT >> swapP) .* id -- (Nat + U) * (Nat' * a) (sym foldswapP) ...

We can then distribute out the Nat + U again:

... $ do trace id -- (Nat' * (Nat * a)) + (U * (Nat * a)) -- (Nat' + U) * (Nat * a) sym distrib >> fold) .* id -- Nat' * (Nat * a) (swapPfold) -- Nat * (Nat' * a) swapBacT >> swapP) .* id -- (Nat + U) * (Nat' * a) (sym foldswapP) -- (Nat * (Nat' * a)) + (U * (Nat' * a)) distrib ...

And finally, we apply our step iso to the internal state (we do this after the distrib so that we don’t apply the combinator if the incoming number was 0). The fruits of our labor are presented in entirety:

iterNat :: (a <=> a) -> ( Nat * a <=> Nat * a) (aa)a) = do iterNat step sym unite $ do trace id -- (Nat' * (Nat * a)) + (U * (Nat * a)) -- (Nat' + U) * (Nat * a) sym distrib >> fold) .* id -- Nat' * (Nat * a) (swapPfold) -- Nat * (Nat' * a) swapBacT >> swapP) .* id -- (Nat + U) * (Nat' * a) (sym foldswapP) -- (Nat * (Nat' * a)) + (U * (Nat' * a)) distrib ( id .* ( id .* step)) .+ id -- (Nat * (Nat' * a)) + (U * (Nat' * a)) step)) .+ id -- (Nat' * (Nat * a)) + (U * (Nat' * a)) swapBacT unite

Lo and behold, the types now line up, and thus quod erat demonstrandum. The implementation of isEven is now trivial:

isEven :: Nat * Bool <=> Nat * Bool = iterNat not isEveniterNat

which computes if a Nat is even in the case the incoming Bool is false , and whether it is odd otherwise.

Lists

James and Sabry provide a sketch of how to define lists, but I wanted to flesh out the implementation to test my understanding.

For reasons I don’t pretend to understand, Haskell won’t let us partially apply a type synonym, so we’re forced to write a higher-kinded data definition in order to describe the shape of a list.

-- To be read as @type ListF a b = U + (a * b)@. data ListF a b a b = Nil | Cons a b a b

We can then get the fixpoint of this in order to derive a real list:

type List a = Fix ( ListF a) a)

And to get around the fact that we had to introduce a wrapper datatype in order to embed this into Haskell, we then provide an eliminator to perform “pattern matching” on a List a . In a perfect world, this function would just be sym fold , but alas, we must work with what we have.

liste :: List a <=> U + (a * List a) (aa) = Iso to from listeto from where Fix Nil ) = InL U to ( Fix ( Cons a b)) = InR ( Pair a b) to (a b))a b) InL U ) = Fix Nil from ( InR ( Pair a b)) = Fix ( Cons a b) from (a b))a b)

From here, it is trivial to write cons :

cons :: a * List a <=> List a = do cons -- U + (a * List a) just -- List sym liste

However, introducing a list via nil is actually quite complicated. Note the parallels with the natural numbers, where it was trivial to define succ but required a clever trick to introduce a zero .

We begin with a lemma that moves a coproduct:

swapCbaP :: (a + b) + c <=> (c + b) + a (ab)(cb) = do swapCbaP -- a + (b + c) sym assocP -- (b + c) + a swapP .+ id -- (c + b) + a swapP

And given that, we can write an isomorphism between any a and any b . The catch, of course, is that you can never witness such a thing since it obviously doesn’t exist. Nevertheless, we can use it to convince the type checker that we’re doing the right thing in cases that would diverge in any case.

diverge :: a <=> b = trace $ do divergetrace id -- (a + b) + a .+ id -- (b + a) + a swapP -- (a + a) + b swapCbaP .+ id -- a + b sym injectR -- b + a swapP .+ id -- (b + b) + a right -- (a + b) + b swapCbaP

Finally we can implement nil using the same trick we did for zero – use trace to vacuously introduce exactly the type we need, rip out the result, and then divergently reconstruct the type that trace expects.

nil :: U <=> List a = trace $ do niltrace id -- (a * List a) + U -- U + (a * List a) swapP -- List a sym liste -- U * List a sym unite .* id -- (U + U) * List a just -- (U * List a) + (U * List a) distrib .* id ) .+ unite -- (a * List a) + List a (divergeunite

Induction on Lists

In a manner spiritually similar to iterNat , we can define iterList :: (a * z <=> b * z) -> (List a * z <=> List b * z) . The semantics are mostly what you’d expect from its type, except that the resulting List b is in reverse order due to having to be constructed as the List a was being destructed. We present the implementation here for completeness but without further commentary.

iterList :: (a * z <=> b * z) -> ( List a * z <=> List b * z) (az)z) = do iterList f sym unite $ do trace -- ((b * List b) * (List a * z)) + (U * (List a * z)) -- ((b * List b) + U) * (List a * z) sym distrib >> sym liste) .* id -- List b * (List a * z) (swapPsym liste) -- List a * (List b * z) swapBacT .* id -- (U + (a * List a)) * (List b * z) liste -- (U * (List b * z)) + ((a * List a) * (List b * z)) distrib ( .+ ) id $ -- (U * (List b * z)) + ... do .* id -- ((List a * a) * (List b * z)) swapT -- ((List a * List b) * (a * z)) swapAcbdT id .* f -- ((List a * List b) * (b * z)) -- ((List a * b) * (List b * z)) swapAcbdT -- ((List a * b) * (List b * z)) + (U * (List b * z)) swapP .* id ) .+ id -- ((b * List a) * (List b * z)) + (U * (List b * z)) (swapT .+ id -- ((b * List b) * (List a * z)) + (U * (List b * z)) swapAcbdT unite swapAcbdT :: (a * b) * (c * d) <=> (a * c) * (b * d) (ab)(cd)(ac)(bd) = do swapAcbdT sym assocT id .* sw sw assocT

From here, the functional programming favorite map is easily defined:

map :: (a <=> b) -> ( List a <=> List b) (ab)b) map f = do sym unite swapT $ f .* id -- map iterList id -- reverse to original order iterList swapT unite

Remnants

The bulk of the remainder of the paper is an extension to the reversible semantics above, introducing create :: U ~> a and erase :: a ~> U where (~>) is a non-reversible arrow. We are shown how traditional non-reversible languages can be transformed into the (~>) -language.

Of more interest is James and Sabry’s construction which in general transforms (~>) (a non-reversible language) into (<=>) (a reversible one). But how can such a thing be possible? Obviously there is a trick!

The trick is this: given a ~> b , we can build h * a <=> g * b where h is “heap” space, and g is “garbage”. Our non-reversible functions create and erase thus become reversible functions which move data from the heap and to the garbage respectively.

Unfortunately, this is a difficult thing to model in Haskell, since the construction requires h and g to vary based on the axioms used. Such a thing requires dependent types, which, while possible, is quite an unpleasant undertaking. Trust me, I actually tried it.

However, just because it’s hard to model entirely in Haskell doesn’t mean we can’t discuss it. We can start with the construction of (~>) :

{-# LANGUAGE GADTs #-} data a ~> b where Arr :: (a <=> b) -> (a ~> b) (ab)(ab) Compose :: (a ~> b) -> (b ~> c) -> (a ~> c) (ab)(bc)(ac) First :: (a ~> b) -> (a * c ~> b * c) (ab)(ac) Left :: (a ~> b) -> (a + c ~> b + c) (ab)(ac) Create :: U ~> a Erase :: a ~> U

The axioms here are quite explanatory and will not be discussed further. A point of note, however, is that Arr allows arbitrary embeddings of our iso (<=>) language in this arrow language.

The semantics of Create is given by induction:

\[

ewcommand{\u}{\text{U}} \begin{align*} \text{create U} & \mapsto \u \\ \text{create}(a + b) & \mapsto \text{InL } (\text{create } a) \\ \text{create}(a \times b) & \mapsto (\text{create } a, \text{create } b) \end{align*} \]

With the ability to create and erase information, we’re (thankfully) now able to write some everyday functions that you never knew you missed until trying to program in the iso language without them. James et al. give us what we want:

fstA :: a * b ~> a = do fstA -- b * a arr swapT -- U * a first erase -- a arr unite

In addition to projections, we also get injections:

leftA :: a ~> a + b = do leftA $ sym unite -- U * a arrsym unite -- (a + b) * a first create -- (a + b) * a arr leftSwap -- a + b fstA leftSwap :: (a + b) * a <=> (a + b) * a (ab)(ab) = do leftSwap -- (a * a') + (b * a') distrib .+ id -- (a' * a) + (b * a') swapT -- (a' + b) * a sym distrib

And the ability to extract from a coproduct:

join :: a + a ~> a = do join $ do arr .+ sym unite -- (U * a) + (U * a) sym unitesym unite -- (U + U) * a sym distrib -- a * (U + U) swapT -- a fstA

We are also provided with the ability to clone a piece of information, given by structural induction. Cloning U is trivial, and cloning a pair is just cloning its projections and then shuffling them into place. The construction of cloning a coproduct, however, is more involved:

clone :: a + b ~> (a + b) * (a + b) (ab)(ab) = do clone $ do left -- (a * a) + b clone -- ((a + b) * a) + b first leftA -- (a * (a + b)) + b arr swapT -- b + (a * (a + b)) arr swapP $ do left -- (b * b) + (a * (a + b)) clone -- ((b + a) * b) + (a * (a + b)) first leftA -- (b * (b + a)) + (a * (a + b)) arr swapT $ do arr -- (a * (a + b)) + (b * (b + a)) swapP id .+ ( id .* swapP) -- (a * (a + b)) + (b * (a + b)) swapP) -- (a + b) * (a + b) sym distrib

It should be quite clear that this arrow language of ours is now more-or-less equivalent to some hypothetical first-order version of Haskell (like Elm?). As witnessed above, information is no longer a linear commodity. A motivated programmer could likely get work done in a 9 to 5 with what we’ve built so far. It probably wouldn’t be a lot of fun, but it’s higher level than C at the very least.

The coup de grace of Information Effects is its construction lifting our arrow language back into the isomorphism language. The trick is to carefully construct heap and garbage types to correspond exactly with what our program needs to create and erase. We can investigate this further by case analysis on the constructors of our arrow type:

Arr :: (a <=> b) -> (a ~> b) (ab)(ab)

As we’d expect, an embedding of an isomorphism in the arrow language is already reversible. However, because we need to introduce a heap and garbage anyway, we’ll use unit.

Since we can’t express the typing judgment in Haskell, we’ll use a sequent instead:

\[

ewcommand{\lifted}[3]{\text{lift } #1 : #2 \leftrightarrow #3}

ewcommand{\arr}{\rightsquigarrow} \frac{\text{arr } f : a \arr b}{\lifted{(\text{arr } f)}{\u \times a}{\u \times b}} \]

Assuming we have a way of describing this type in Haskell, all that’s left is to implement the lift ing of our iso into the enriched iso language:

Arr f) = id .* f lift (f)

Compose :: (a ~> b) -> (b ~> c) -> (a ~> c) (ab)(bc)(ac)

Composition of arrows proceeds likewise in a rather uninteresting manner. Here, we have two pairs of heaps and garbages, results from lifting each of the arrows we’d like to compose. Because composition will run both of our arrows, we’ll need both heaps and garbages in order to implement the result. By this logic, the resulting heap and garbage types are pairs of the incoming ones.

\[ \frac{\lifted{f}{h_1\times a}{g_1\times b},\; \lifted{g}{h_2\times b}{g_2\times c}}{\lifted{(g \circ f)}{(h_1\times h_2)\times a}{(g_1\times g_2)\times c}} \]

We can express the resulting combinator in Haskell:

Compose f g) = do lift (f g) id -- (H1 * H2) * a .* id -- (H2 * H1) * a swapT -- H2 * (H1 * a) sym assocT id .* lift f -- H2 * (G1 * b) lift f -- (H2 * G1) * b assocT .* id -- (G1 * H2) * b swapT -- G1 * (H2 * b) sym assocT id .* lift g -- G1 * (G2 * c) lift g -- (G1 * G2) * c assocT

First :: (a ~> b) -> (a * c ~> b * c) (ab)(ac)

Lifting arrows over products again is uninteresting – since we’re doing nothing with the second projection, the only heap and garbage we have to work with are those resulting from the lifting of our arrow over the first projection.

\[ \frac{\lifted{f}{h\times a}{g\times b}} {\lifted{(\text{First } f)}{h\times (a\times c)}{g\times (b\times c)}} \]

In Haskell, our resulting combinator looks like this:

First f) = do lift (f) id -- H * (a * c) -- (H * a) * c assocT f .* id -- (G * b) * c -- G * (b * c) sym assocT

Left :: (a ~> b) -> (a + c ~> b + c) (ab)(ac)

Finally, we get to an interesting case. In the execution of Left , we may or may not use the incoming heap. We also need a means of creating a b + c given a b or given a c . Recall that in our iso language, we do not have create (nor relatedly, leftA ) at our disposal, and so this is a harrier problem than it sounds at first.

We can solve this problem by requiring both a b + c and a c + b from the heap. Remember that the Toffoli construction (what we’re implementing here) will create a reversible gate with additional inputs and outputs that gives the same result when all of its inputs have their default values (ie. the same as those provided by create ’s semantics). This means that our incoming b + c and c + b will both be constructed with InL .

Given this, we can thus perform case analysis on the incoming a + c , and then use leftSwap from earlier to move the resulting values into their coproduct.

What does the garbage situation look like? In the case we had an incoming InL , we will have used up our function’s heap, as well as our b + c , releasing the g , b (the default value swapped out of our incoming b + c ), and the unused c + b .

If an InR was input to our isomorphism, we instead emit the function’s heap h , the unused b + c , and the default c originally in the heap’s coproduct.

Our final typing judgment thus looks like this:

\[ \frac{\lifted{f}{h\times a}{g\times b}}{\lifted{(\text{Left f})}{h'\times (a + c)}{g' \times (b + c)}} \]

\[ h' = h\times ((b + c) \times(c + b)) \\ g' = (g\times (b\times(c + b))) + (h\times ((b+c)\times c)) \]

and is rather horrifyingly implemented:

Left f) = do lift (f) swapT distrib .+ rightSide leftSide frightSide sym distrib leftSide :: (h * a <=> g * b) (hb) -> (a * (h * ((b + c) * (c + b))) <=> (g * (b * (c + b))) * (b + c)) (a(h((bc)(cb)))(g(b(cb)))(bc)) = do leftSide f -- (H * ((b + c) * (c + b))) * a swapT .* id -- (((b + c) * (c + b)) * H) * a swapT -- ((b + c) * (c + b)) * (H * a) sym assocT id .* f -- ((b + c) * (c + b)) * (G * b') .* id -- ((c + b) * (b + c)) * (G * b') swapT -- ((c + b) * G) * ((b + c) * b') sw2 id .* leftSwap -- ((c + b) * G) * ((b' + c) * b) leftSwap .* swapT -- (G * (c + b)) * (b * (b' + c)) swapTswapT -- ((G * (c + b)) * b) * (b' + c) assocT .* id -- (G * ((c + b) * b)) * (b' + c) sym assocT ( id .* swapT) .* id -- (G * (b * (c + b))) * (b' + c) swapT) rightSide :: c * (h * ((b + c) * (c + b))) <=> (h * ((b + c) * c)) * (b + c) (h((bc)(cb)))(h((bc)c))(bc) = do rightSide -- c' * (H * ((b + c) * (c + b))) swapT .* id -- ((H * (b + c)) * (c + b)) * c' assocT -- (H * (b + c)) * ((c + b) * c') sym assocT id .* leftSwap -- (H * (b + c)) * ((c' + b) * c) leftSwap id .* swapT -- (H * (b + c)) * (c * (c' + b)) swapT -- ((H * (b + c)) * c) * (c' + b) assocT .* swapP -- (H * ((b + c) * c)) * (b + c') sym assocTswapP

The home stretch is within sight. We have only two constructors of our arrow language left. We look first at Create :

Create :: U ~> a

Because we’ve done all of this work to thread through a heap in order to give us the ability to create values, the typing judgment should come as no surprise:

\[ \frac{}{\lifted{\text{create}}{a\times\u}{\u\times a}} \]

Our heap contains the a we want, and we drop our incoming U as garbage. The implementation of this is obvious:

Create = swapT liftswapT

We’re left with Erase , whose type looks suspiciously like running Create in reverse:

Erase :: a ~> U

This is no coincidence; the two operations are duals of one another.

\[ \frac{}{\lifted{\text{erase}}{\u\times a}{a\times\u}} \]

As expected, the implementation is the same as Create :

Erase = swapT liftswapT

And we’re done! We’ve now constructed a means of transforming any non-reversible program into a reversible one. Success!

Summary

Still here? We’ve come a long way, which we’ll briefly summarize. In this paper, James and Sabry have taken us through the construction of a reversible language, given a proof that it’s Turing-complete, and given us some simple constructions on it. We set out on our own to implement lists and derived map for them.

We then constructed a non-reversible language (due to its capability to create and erase information), and then gave a transformation from this language to our earlier reversible language – showing that non-reversible computing is a special case of its reversible counterpart.

Information Effects ends with a short discussion of potential applications, which won’t be replicated here.

Assuming I understand the physics correctly (which I probably don’t), the fact that these reversible functions do not increase entropy implies that they should be capable of shunting information for near-zero energy. Landauer’s Principle and Szilard’s engine suggests that information entropy and thermodynamic entropy are one and the same; if we don’t increase entropy in our computation of a function, there is nowhere for us to have created any heat.

That’s pretty remarkable, if you ask me. Together with our construction from any non-reversible program to a reversible one, it suggests we should be able to cut down on our CPU power usage by a significant order of magnitudes.

An obvious limitation of what we’ve built here today is that it is first-order, which is to say that functions are not a first class citizen. I can think of no immediate problem with representing reversible functions in this manner. We’d need to move our (<=>) directly into the language.

id would provide introduction of this type, and (>>) (transitivity) would allow us to create useful values of the type. We’d also need a new axiom:

apply :: a * (a <=> b) <=> b * (b <=> a) (ab)(ba)

which would allow us to use our functions. We should also expect the following theorems (which may or may not be axioms) due to our iso language forming a cartesian closed category:

product :: (a <=> (b * c)) <=> (a <=> b) * (a <=> c) (a(bc))(ab)(ac)

Things that we’d expect to be theorems but are not are:

terminal :: U <=> (a <=> U ) (a select :: a <=> ( U <=> a) a) coproduct :: (a <=> (b + c)) <=> (a <=> b) + (a <=> c) (a(bc))(ab)(ac)