Logic, Thought, and Action 21

Chapter 21

ALGORITHMS FOR RELEVANT LOGIC

Paul Gochet, Pascal Gribomont and Didier Rossetto∗Universite de Li`´ ege

Abstract The classical analytic tableau method has been extended successfully tomodal logics (see e.g. Fitting 1983; Fitting 1993; Gore 1992) and also to´relevant and paraconsistent logics Bloesch 1993a; Bloesch 1993b. Theclassical connection method has been extended to modal and intuition-istic logics Wallen 1990, and the purpose of this paper is to investigatewhether a similar adaptation to relevant logic is possible. A hybridmethod is developed for B+, with a specific solution to the “multiplicityproblem”, as in the technique of modal semantic diagrams introducedin Hughes and Cresswell 1968. Proofs of soundness and completenessare also given.

1. IntroductionThe sequent calculi and tableau methods suffer from three kinds of

redundancies : duplication of redundant information, the considerationof reductions that do not advance the search toward finding a proof, andthe need to distinguish derivations that differ in the order in which se-quent rules are applied Wallen 1990, p. 82. The connection method hasproved computationally more efficient for classical, modal and intuition-

∗The authors thank Prof. Jean-Paul van Bendegem, editor of ”Logique et Analyse”, forpermission to reprint “Algorithms for Relevant Logic” which was published in Logique etAnalyse, 150-151-152 Special issue dedicated to the memory of Leo Apostel, pp.329-346,´1995.We are grateful to the research team of the Automated Reasoning Project of the AustralianNational University (Canberra) for their invaluable help. We are very grateful to Dr. Bloeschfor granting us the permission to quote his Ph.D. thesis. We received substantial help fromDr. Gore and Dr. Lugardon. We also thank Dr. van der Does and Dr. Herzig. This work was´supported by a grant of the F.N.R.S. (project 8.4536.94) and an earlier version was discussedat the Centenary Conference of the Lvov-Warsaw School of Logic (Nov. 1995). We expressour gratitude to our sponsors and our hosts, and to Dr. Gailly and Mrs Gailly Goffaux forseveral corrections. We owe the topic of this paper to the late Prof. Sylvan.

D. Vanderveken (ed.), Logic, Thought & Action, 479-496.©c All rights reserved. Printed by Springer, The Netherlands.2005

480 LOGIC, THOUGHT AND ACTION

istic logics; it may therefore be useful to extend it to other non-classicallogics. This is especially true for logics used in artificial intelligence,for which efficient theorem proving techniques are needed. Bloesch hasconjectured that such an extension is feasible in several cases Bloesch1993a, p. 24 :

“It is not clear how to create connection method style proof systems formany non classical logics. In particular the relevant and paraconsistentlogics seem particularly difficult since the law of noncontradiction doesnot hold in most paraconsistent logics and many relevant logics. By re-lying heavily on the properties of classical logic, the connection methodgains great efficiency but at the cost of poor flexibility. It does howeverseem fair to conjecture that techniques, such as using truth signs torepresent exclusive semantic classes, developed in later chapters, couldbe applied to the connection method.”

The main purpose of this paper is to investigate the connection methodstyle for B+, the most basic system of relevant logic, with the simpli-fied semantics introduced in Priest and Sylvan 1992; Restall 1993. Theconnection method presented here draws from two sources : Wallen’sconnection method for modal logic K and Bloesch’s tableau methodfor relevant and paraconsistent logics. Both methods will have to bemodified to fit our purpose. B+ is known to be decidable. We presenta decision procedure which automatically supplies finite models whenapplied to a formula which happens to be satisfiable.

This paper goes on as follows : the axiomatic system B+ is recalled,Bloesch’s Tableau Method for B+ is briefly presented, Wallen’s Connec-tion Method for modal logic is adapted to B+, and the soundness andthe completeness of the extension are proven.

2. The axiomatic method for relevant logic B+

2.1 Axioms(i) A→ A

(ii) A→ (A ∨B), B → (A ∨B)

(iii) (A ∧B)→ A, (A ∧B)→ B

(iv) (A ∧ (B ∨ C))→ ((A ∧B) ∨ C)

(v) ((A→ B) ∧ (A→ C))→ (A→ (B ∧ C))

(vi) ((A→ B) ∧ (B → C))→ ((A ∨B)→ C)

Algorithms for Relevant Logic 481

2.2 Rules(i) A , A→B

B (modus ponens), and its disjunctive form.

(ii) A , BA∧B (adjunction) and its disjunctive form

(iii) A→B , C→D(B→C)→ (A→D) (affixing rule) and its disjunctive form

Comment. If A1,...,An

B is a rule then its disjunctive form is the ruleC∨A1,...,C∨An

C∨B .

2.3 ExampleLet us prove that formula 21.1 is a theorem of B+.

(p→ q)→ ((s→ q) ∨ ((r ∧ p)→ q)) (21.1)

Proof.

(i) (r ∧ p)→ p (axiom 3)

(ii) q → q (axiom 1)

(iii) (p→ q)→ ((r ∧ p)→ q) (rule 3 (1, 2))

(iv) (p→ q)→ (p→ q) (axiom 1)

(v) ((r ∧ p)→ q)→ ((s→ q) ∨ ((r ∧ p)→ q)) (axiom 2)

(vi) ((p→ q)→ ((r∧p)→ q))→ ((p→ q)→ ((s→ q)∨((r∧p)→ q)))(rule 3 (4, 5))

(vii) (p→ q)→ ((s→ q) ∨ ((r ∧ p)→ q)) (rule 1 (3, 6))

This theorem will be proven again by the method presented here.

3. Tableau method for relevant logic B+

The principle of the semantic tableau method is rather straightfor-ward. It is a systematic search for a model that falsifies the formulabeing checked for validity. If such a model does not exist, the formula isvalid.

The formula ϕ to be proven is assumed first to be false in world w0

(we write Fw0ϕ). Next it is reduced by applying rules which remove theconnectives stepwise starting with the less deeply nested connective untileach branch of the tableau becomes closed or gives rise to a model of¬ϕ. If all the branches of the tree are closed (i.e. there is no model that


falsifies ϕ), Fw0ϕ cannot be forced and ϕ is valid. This process is notdeterministic. Generally several tableaux can be produced dependingon which order we chose when we apply the reduction rules.

The reduction rules are based on the truth conditions. For example,A∧B is true iff A and B are true. The truth conditions of the relevantconditional bear some resemblance with those of the classical conditional.Indeed, when the relevant conditional A→ B is false the antecedent Ais true and the consequent B is false, and when the relevant conditionalis true, the antecedent is false or the consequent is true. However, as therelevant conditional is intensional, something more is needed to captureits meaning. Just like modal formulas, the intensional formula A → Bis evaluated at a world w and an accessibility relation relates the worldw to the worlds w′ and w′′ at which A and B respectively are evaluated.The interpretation of relevant implication involves three worlds, insteadof two for modal operators, so the relevant accessibility relation R isternary.1 The contrast between FwA → B and TwA → B matches thecontrast between FwA and TwA. Formula A → B is false at w iffw′ and w′′ exist such that Rww′w′′ , A is true at w′ and B is false at w′′,just as A is false at w iff there exists a world w′ such that Rww′ andA is false at w′. Formula A→ B is true iff for all w′ and w′′ such thatRww′w′′ , A is false at w′ or B is true at w′′.

Taking advantage of this semantics, Bloesch Bloesch 1993a, p. 88 givesreduction rules for FwA→ B and TwA→ B. Whenever a formula suchas FwA → B appears on a branch, we record, for that branch, thatRww′w′′ and we add the formulas Tw′A and Fw′′B to the branch, wherew′ and w′′ are new worlds, i.e. they do not appear on the branch, andw′ = w′′ iff w = w0 (i.e. the base world). Whenever a formula such asTwA → B appears on a branch, if we have Rww′w′′ on that branch, wemay create two subbranches with the formulas Fw′A on one and Tw′′Bon the other.

Since the tableau rule for F→ is a hybrid rule which bears similarityboth to π-rule (modal rule dealing with T♦) and to α-rule, it is appro-priate to introduce the category of “πα-rule”. Analogously a tableaurule for T→ is a hybrid rule which falls under the heading of “νβ-rule”(modal ν-rule deals with T).

The tableau rules based on the semantics of the connectives fall there-fore in four types : α, β, πα, and νβ (figure 1).

1A frame in modal logic is a graph; a frame in relevant logic is a hypergraph.


α α1 α2

Twix ∧ y Twix TwiyFwix ∨ y Fwix Fwiy

πα πα1 πα2

Fwix → y Twj x Fwky

wj and wk (j = k iff i = 0) are newworlds and Rwiwjwk is added

β β1 β2

Fwix ∧ y Fwix FwiyTwix ∨ y Twix Twiy

νβ νβ1 νβ2

Twix → y Fwj x Twky

wj and wk are such that Rwiwjwk waspreviously added

Figure 1. Reduction rules for B+

(i) Prolongation rules (or α-rules). Each α-node gives rise to a pro-longation of the current branch. Two successive nodes are added,one of which is labelled with α1 and the other is labelled with α2.

(ii) Branching rules (or β-rules). Each β-node splits the branch intotwo subbranches one of which bears a node labelled with β1 andthe other a node labelled with β2.

(iii) Relevant prolongation rule (or πα-rule). Each πα-node gives riseto a prolongation of the branch. Two successive nodes are added,one of which is labelled with πα1 and the other is labelled withπα2.

(iv) Relevant branching rule (or νβ-rule). Each νβ-node splits thebranch into two subbranches one of which bears a node labelledwith νβ1 and the other a node labelled with νβ2. The νβ-rule canbe used µ times; µ, also called the multiplicity of the νβ-formula,is the number of pairs of worlds related to wi previously introducedby the application of a πα-rule (Rwiwjwk

was previously added tothe branch).

Comment. The parameter µ associated with a νβ-formula has a definitevalue only at the end of the derivation. Here is the feature that will raisethe multiplicity problem in the connection method.

4. Connection method extended to relevant logicThe principle of the connection method is the same as that of the

tableau method : a formula is said to be proven by a connection proofwhenever its attempted refutation fails. The connection method pro-ceeds in three stages : a syntactic tree, an indexed formula tree and apath tree are successively built.


4.1 The syntactic treeThe syntactic tree or formation tree displays the structure of the for-

mula. For example, figure 2 shows the syntactic tree of formula 21.1.The nodes are numbered by traversing the tree depthfirst, from left toright.

Figure 2. Syntactic Tree

4.2 The indexed formula treeIn tableau proof trees, a derivation is done by removing connectives in

accordance with logical rules of elimination which we call reduction steps.These steps are graphically depicted (e.g. the elimination of an assertedconjunction gives rise to a prolongation of the branch, the elimination ofa denied conjunction gives rise to the splitting of the branch,. . . ). Thesign of the initial formula tested for inconsistency is F since a refutationproof is intended. The sign of the other formulas is determined by thesign of the input formula and by the reduction rule applied to it (e.g.the denial of a formula whose main connective is a conjunction leads tothe disjunction of the denials of each conjunct).

In the connection method, these pieces of information are encodedin the indexed formula tree. The indexed formula tree associated witha signed formula ϕ can be represented as an array (figure 3 shows theindexed formula tree of formula 21.1). Each line of this array is a recordrepresenting a signed subformula of ϕ obtained when applying recur-sively the reduction rules (figure 1) to ϕ.

The indexed formula tree has seven columns. The first column is akey identifying the line;2 the key will be ai, where i is the index of the

2The path tree will deal with these indices to represent the (sub)formulas in an efficient way.


k pol(k) lab(k) Pt(k) St(k) w(k) h(k)

a0 F (p→q)→((s→q)∨((r∧p)→q)) πα − w0 0a1 T p → q νβ πα1 w1 1a12 F p − νβ1 w2 4

a13 T q − νβ2 w3 4

a22 F p − νβ1 w4 7

a23 T q − νβ2 w5 7

a4 F (s → q) ∨ ((r ∧ p) → q) α πα2 w1 1a5 F s → q πα α1 w1 2a6 T s − πα1 w2 3a7 F q − πα2 w3 3a8 F (r ∧ p) → q πα α2 w1 2a9 T r ∧ p α πα1 w4 5a10 T r − α1 w4 6a11 T p − α2 w4 6a12 F q − πα2 w5 5

Figure 3. Indexed Formula Tree

subformula in the syntactic tree of ϕ. Some formulas have to be repeated(see later); superscripts are used to distinguish the multiple occurrencesof these formulas, and are transmitted to their children. The secondcolumn, named pol(k) (where k is the key), records the polarity of theformula (T or F). The third column, named lab(k), records the label,i.e. the signed subformula itself. Each non atomic signed formula hasa type, which can be α, β, πα or νβ, and also two components, calledchildren, as we saw in figure 1. Each child has its own type, called theprimary type; its secondary type is the (primary) type of the parentformula, with the subscript 1 or 2, as indicated in figure 1. The fourthcolumn, named Pt(k), records the (primary) type of the subformula; anatomic formula has no primary type (− in the fourth column). The fifthcolumn, named St(k), records the secondary type of the subformula.The main formula ϕ has no secondary type (− in the fifth column). Thecolumns “Pt(k)” and “St(k)” are introduced merely to support the un-derstanding of the indexed formula tree. In relevant logic, as in modallogic, the polarity is ascribed to a formula relatively to a given world.The sixth column, named w(k), records this piece of information; worldsare given names from the unlimited sequence w0, w1, w2, . . . Indeed twoidentical formulas of opposite signs such as TA and FA will combinetogether to yield a contradiction only if they are evaluated at the sameworld. The last column, named h(k), records the history : it containsthe number of the step during which the line is created. Column 7 is also


introduced to support the understanding of the indexed formula tree.

We now give the iterative procedure to construct the indexed formulatree. Throughout the execution, the procedure maintains for each line abinary variable, whose value can be “active” or “passive”. Furthermore,for each line of type νβ, a set Ω of ordered pairs of worlds is maintained.

Each step of the procedure selects an active line and generates twoor more new lines 1, 2, . . . In every case, the label of a new line is achild of the label of the parent line, according to the type-dependentreduction rules given in figure 1. After the generation, line becomespassive, whereas lines 1, 2, . . . are active if their label is non atomic andpassive otherwise. The set Ω associated with a new νβ-line is empty.

Initial step 0. Create an initial line.

- The initial line has key a0; its label is the given signed formula,to which we ascribe the initial world w0. The history is 0. Theline is made active if its (primary) type is α, β or πα, and passiveotherwise.

Iterative step n > 0. Select an active line of key k.

- If Pt(k) = α, two lines are added to the indexed formula tree.Their labels respectively are the α1-child and the α2-child of lab(k)determined by the α-reduction rule; their indices are extractedfrom the syntactic tree; their world and history are w(k) and n,respectively.

- If Pt(k) = β, two lines are added to the indexed formula tree.Their labels respectively are the β1-child and the β2-child of lab(k),determined by the β-reduction rule; their indices are extractedfrom the syntactic tree and their world is w(k). The history ofboth new lines is n.

- If Pt(k) = πα, two lines are added to the indexed formula tree.Their labels respectively are the πα1-child and the πα2-child oflab(k), determined by the πα-reduction rule; their indices are ex-tracted from the syntactic tree. The worlds wj and wk associatedwith these two formulas must not have been used before, say thefirst elements of the world sequence which differ from all worldsassociated with existing lines. The history of both new lines is n.Furthermore, each existing νβ-line k′ (i.e. Pt(k′) = νβ) such thatw(k′) = w(k) is made active again.


- If Pt(k) = νβ, the set E of all passive existing πα-lines k′ (i.e.Pt(k′) = πα) such that w(k′) = w(k) is determined. For eachk′ ∈ E we do the following. As k′ is passive, it has two children,say k′′ and k′′′, whose associated worlds are w(k′′) and w(k′′′). If(w(k′′), w(k′′′)) is not an element of the set Ω(k) associated withline k, two new lines, say and ′, are added to the indexed for-mula tree. Signed subformula lab() and lab(′) are the children oflab(k), as determined by the νβ-reduction rule; the indices and′ are extracted from the syntactic tree. Remember that a super-script is used to distinguish the different pairs of children of theparent formula. The associated worlds w() and w(′) are w(k′′)and w(k′′′) respectively, and the ordered pair (w(k′′), w(k′′′)) isadded to Ω(k). The history of each new line is n.Comment. It should be emphasized that a νβ-line k may switchseveral times from active to passive and conversely during the ex-ecution. Besides, each step about the νβ-line k may introduce anynumber of new ordered pairs of children for k, up to the size ofthe set E. Last, the worlds associated with the children are notinherited from the parent line k, but from other lines, of secondarytype παi (i = 1, 2).

4.3 The path treeThe path tree (it is actually an acyclic graph, see Gribomont and

Rossetto 1995 for an example) has the same role as a tableau proof tree,but its construction is computationally more efficient. The path treecomes from the following recursive definition of a path.

Basis.

- If a0 is the root of the indexed formula tree, a0 is a path.

Recursion.

- If S is a path containing the node α,3

(S \ α) ∪ α1 ∪ α2 is a path.

- If S is a path containing the node β,(S \ β) ∪ β1 and (S \ β) ∪ β2 are paths.

- If S is a path containing the node πα,(S \ πα) ∪ πα1 ∪ πα2 is a path.

- If S is a path containing the node νβ,S ∪ t1, . . . , tµ, where ti (1 ≤ i ≤ µ) is either νβi

1 or νβi2, are

paths.

3I.e. a node of the indexed formula tree whose label is a signed formula of primary type α.


Comments. There are 2µ such paths, where µ is the multiplicityassociated with νβ. The µ ordered pairs (νβ1, νβ2) used hereare the µ ordered pairs (among the ordered pairs (νβ1, νβ2) gotduring the building of the indexed formula tree) which are framecompatible with the nodes of the current path, i.e. νβ1 and νβ2

are associated with worlds already involved in the current path.4

Each step in the construction of the path tree consists of the applicationof a rule (α, β, πα, or νβ) to a formula. These applications generate newpaths. It should be emphasized here that as soon as a the successorsof a path have been determined, this path can be erased. In fact itis erased when the connection method is implemented on a computer;the path tree never resides wholly in the computer memory. Only itsleaves, i.e., the atomic paths, are saved and determine models of ¬ϕ ifthey do not contain connections. The atomic paths contain only atomicformulas and (vacuously true) νβ-formulas.

The formula tested is a theorem if and only if each leaf of the path treecontains a connection, that is, two signed formulas TwiA and Fwj B,with wi = wj and A identical to B.

To ensure soundness, the path tree has to respect a reduction orderingwhich combines two orderings : the subformula ordering and the modalordering. The reduction ordering (denoted ) is the transitive closureof the union of the subformula ordering (denoted ) and the modalordering (denoted M), so = ( ∪ M)∗ Wallen 1990.

Therefore the recursion has to be applied in the following way.

(i) Respecting the order of the world indices : consider nodes relatedto world wi only when all nodes attached to world wj with j < ihave been considered.

(ii) Respecting the necessity order : consider νβ-nodes related to worldwi only when all other nodes related to world wi have been con-sidered.

This is automatically done if the recursion is applied in the order recor-ded in the seventh column h(k) of the indexed formula tree.

The reason behind the frame compatibility restriction is this : allowingthe introduction in a path of formulas that are not frame compatible withthe other formulas of the path would amount to allowing a move fromone subbranch to another in a tableau proof tree, which would clearlybe unsound.

Let us observe that if we consider the semantic tableau style of proof,a νβ-formula could introduce µ successive branching that lead to 2µ

splits of the branch. In the path tree, all the 2µ paths generated by

4A syntactic criterion to check frame compatibility between two formulas is the following :formula ϕ is frame compatible with formula ψ if and only if the primary type of their commonancestor in the syntactic tree is neither β nor νβ.


a0a1, a4a1, a5, a8a1, a6, a7, a8

a1, a6, a7, a9, a12

a12, a

22, a6, a1

3, a22, a6, a1

2, a23, a6, a1

3, a23, a6,

a7, a9, a12 a7, a9, a12 a7, a9, a12 a7, a9, a12a1

2, a22, a6, a7, a1

3, a22, a6, a7, a1

2, a23, a6, a7, a1

3, a23, a6, a7,

a10, a11, a12 a10, a11, a12 a10, a11, a12 a10, a11, a12

Figure 4. Path Tree

its instantiation are got in one step. Indeed, in tableau trees we do thebranching and instantiation separately. On the contrary here we combinebranching with modal instantiation when applying νβ-rules. Our policyproduces a reduction of the number of nodes since 2µ nodes are neededinstead of

∑µi=1 2i = 2µ+1 − 2.

When building the path tree, the same operation applies for rules ofany category : we replace the parent formula by its children. Here againour technique diverges from the standard one used in modal logic, wherewhenever a necessity rule is applied, a child is added but the parentformula is maintained; if the nth instantiation has failed to produce acontradiction, a new instantiation takes place; multiplicity is demanddriven Wallen 1990.

Demand driven policy secures completeness. However if the formulatested fails to be unsatisfiable, a loop may occur. The policy advocatedin this paper rules out the risk of non termination as far as B+ is con-cerned. Moreover it automatically supplies finite models whenever theformula tested for inconsistency happens to be satisfiable. Our treatmentof the multiplicity problem establishes a tight correspondence betweenthe νβ-reduction rule and the semantics of νβ-formulas.

The path tree corresponding to formula 21.1 is represented in figure4. To help the reader to understand the path tree, the node developed isunderlined once and the indices standing for connections are underlinedtwice. There are four atomic paths, each of them containing a connec-tion, so no model exists for signed formula Fw0ϕ, and ϕ is a B+-validformula.

We see that the four atomic paths all contain a connection, thereforeno frame that falsifies formula 21.1 exists, i.e. this formula is B+-valid.


5. Soundness and completenessIn this section we prove that our method is a sound and complete

decision procedure for relevant logic B+.

5.1 ComplexityIf A is a (signed or not) formula, its degree d(A) is the number of

internal nodes of the syntactic tree associated with A, i.e. the totalnumber of connectives contained in A. In other words, the degree of anatomic formula is 0; d(¬A) = 1 + d(A) and if is a binary operator,d(AA′) = 1+d(A)+d(A′). The notion of degree also applies to signedformulas : d(TA) = d(FA) = d(A) .

If E is a set of formulas, its complexity c(E) is the sum of the degreesof its elements.

Every finite set of formulas has a finite complexity. Furthermore, withusual notation, we have the following inequalities :

- if E1 ⊆ E2, then c(E1) ≤ c(E2).- c(E ∪ α) > c((E \ α) ∪ α1, α2),

since d(α) > d(α1) + d(α2) .

- c(E ∪ β) > max(c((E \ β) ∪ β1), c((E \ β) ∪ β2)),since d(β) > d(β1) and d(β) > d(β2).

- c(E ∪ πα) > c((E \ πα) ∪ πα1, πα2),since d(πα) > d(πα1) + d(πα2) .

- c(E ∪ νβ) > max(c((E \ νβ) ∪ νβ1), c((E \ νβ) ∪ νβ2)),since d(νβ) > d(νβ1) and d(νβ) > d(νβ2).

5.2 Auxiliary structureAs a preliminary result, we need to know that the execution of the

construction procedure for the indexed formula tree always terminates.We first observe that every step induces the addition of finitely manynew lines (or nodes). This is trivial, even for νβ-steps, since the numberof lines added in a νβ-step cannot exceed 2n where n is the total numberof lines introduced before this step. As a consequence, termination couldbe prevented only if an infinite number of steps could take place.

In order to prove that every execution involves finitely many steps,we will suppose that each step of the construction procedure updatesnot only the indexed formula tree, but also an auxiliary structure de-fined below; it will then be sufficient to prove that this structure canbe updated only finitely many times. The auxiliary structure is a set of


hypertrees; the nodes of the hypertrees are worlds (so each hypertree isa frame). Each node wi is labelled with a set (wi) of signed formulas.

We describe now the update induced on the auxiliary structure byeach step of the construction procedure, for a signed formula ϕ.Initial step 0. (Creation of the initial line.)

- The structure initially contains only one frame; this frame consistsof a single node w0, whose label contains a single element, whichis the initial signed formula : (w0) = ϕ.

Iterative step n > 0. (Addition of children of line k.)In every case, frames that contain world-node w(k) with signed formula“pol(k) : lab(k)” in (w(k)) already exist in the auxiliary structure.

- If Pt(k) = α, in each frame such that the signed formula α =pol(k):lab(k) is a member of (w(k)), we update (w(k)) by re-placing its element α by the corresponding elements α1 and α2.

- If Pt(k) = β, each frame of the auxiliary structure that has a nodenamed w(k) with the signed formula β = pol(k):lab(k) in its label(w(k)) is updated as follows. First, the frame is replaced by twoidentical frames; second, in the label (w(k)), the element β isreplaced by the corresponding element β1 in the first frame and β2

in the second one.

- If Pt(k) = πα, each frame of the auxiliary structure that has anode named w(k) with the signed formula πα = pol(k):lab(k) inits label (w(k)) is updated as follows. Two nodes w and w′ and thehyperarrow (w(k), w, w′) are added to the frame, where w and w′are the new worlds associated with the children of line k. The newnodes are labelled respectively (w) = πα1 and (w′) = πα2.

- If Pt(k) = νβ, each frame that contains node w(k), with νβ ∈(w(k)) and that has a hyperarrow (w(k), w, w′) for some w and w′is updated as follows, if it has not been done yet. First, the frameis replaced by two identical frames; second, the children signedformulas νβ1 and νβ2 are added to (w) in the first frame and to(w′) in the second frame respectively.

We first observe that duplication of frames is induced by β and νβ-reductions. However, a branching formula can be a subformula of an-other branching formula, so β-reductions and νβ-reductions can be re-used several times. As a result, the number of frame duplication isbounded by d(ϕ)! and the total number of frames in an auxiliary struc-ture cannot exceed 2d(ϕ)!.


Furthermore, each frame of the auxiliary structure is finite. First, sucha frame is a finitary hypertree since the number of successors of any nodeof the frame is bounded by the number of πα-subformulas contained inthe initial formula, and therefore by its degree d(ϕ). Second, the lengthof a hyperbranch cannot exceed d(ϕ) = c(ϕ), since the complexity ofa successor node is strictly less than the complexity of the parent node.An upper bound for the number of nodes in a frame is d(ϕ)d(ϕ)+1, andan upper bound for the total number of nodes in the whole structure istherefore Σϕ =def 2d(ϕ)!d(ϕ)d(ϕ)+1 (this is also an upper bound for thenumber of hyperarrows). Tighter bounds can be found, but we do notneed them here.Comment. The notion of auxiliary structure applies not only to a signedformula, but also to a finite (conjunctive) set of signed formulas.

5.3 Termination proofsWe can now give an upper bound for the length of any execution of the

construction algorithm for the indexed formula tree. Each step inducesat least one of the following operations :

(i) Extension of the auxiliary structure (β-step, πα-step, useful νβ-step)(cannot be performed more than Σϕ times);

(ii) Addition of a formula to a node label (useful νβ-step, πα-step)(cannot be performed more than d(ϕ)× Σϕ times);

(iii) Reduction of an element of a node label (α-step, β-step)(cannot be performed more than d(ϕ)× Σϕ times).

Comment. A νβ-step about line k will do nothing at all if no new or-dered pair of worlds accessible from w(k) has been created since the lastactivation; however, as νβ-lines are “reactivated” by πα-steps, uselesssteps can occur only finitely many times.This completes the termination proof for the indexed formula tree algo-rithm.

In order to prove the termination of the path tree algorithm, we ob-serve that each step adds a finite number of paths to the path tree, whichis a finitary tree. Due to Konig’s lemma, it is sufficient to establish that¨a new path is always strictly less complex than its parent. However, thisis not true for the notion of (multi)-set complexity introduced above;indeed,

c((S \ νβ) ∪ νβ1iββ1, . . . , νβµ

iββµ) < c(S),where ij ∈ 1, 2(1 ≤ j ≤ µ),


cannot be guaranteed.The problem is, we know the multiplicity µ to be finite, but nothing

else. To solve this, we define another well-founded ordering on paths.First, to each path S we associate m(S), a multiset of natural numbers,which are the degrees of the elements of the path. Second, we show that,for some relation , the domain of multisets is well-founded. Third, weshow that, if S′ is produced by the algorithm from the path S, thenm(S′) m(S).

A multiset of natural numbers can be represented as a decreasingsequence of numbers; the lexical ordering on these sequences is definedas follows :

(a1, . . . , an) (b1, . . . , bm)

if there is a number r ∈ N such that r ≤ n, r ≤ m, ai = bi for i = 1, . . . , rand either r = n and r < m, or r < n, r < m and ar+1 < br+1.As (N,<) is a well-ordered set, so is (N∗,). Besides, the lexical ordering induces a (partial) ordering (also noted ) on the set of paths :

S S′ =def m(S) m(S′) ,

and this ordering is well-founded.Last, we observe that each reduction step for a path consists in re-

placing an element of the path by finitely many elements of lower degree;this always leads to a -smaller path.

5.4 Hypertree-modelsThe auxiliary structure is in fact a set of frames, and each frame is an

interpretation of the initial signed formula ϕ, called a hypertree-frame; itis a hypertree-model for ϕ if ϕ holds at the initial world w0. A hypertree-frame is consistent if no world label contains a pair of opposite formulas.We first prove the following lemma.Lemma I. A (hypertree-)frame S is consistent if and only if S, wi |= ψ,for all worlds wi ∈ S and for all signed formulas ψ ∈ (wi).Proof. The “if” part is trivial : if S, wi |= ψ and S, wi |= ξ, then ψ, ξcannot be a pair of opposite signed formulas. For the “only if” part,we proceed by induction on the height of the nodes in the frame. Aleaf w (a node without successor) has height h(w) = 0, and if w isan internal node with successors (w1, w

′1), . . . , (wk, w

′k), then h(w) =

1 + max(h(w1) + h(w′1)), . . . , (h(wk) + h(w′

k)).5Base case. If the world w has no successor in S, the label (w) containsonly atoms and νβ-formulas. The νβ-formulas are vacuously true, and

5Node w has successor (w′, w′′) if (w, w′, w′′) is a hyperarrow.


the atoms are forced to be true at w, which is possible since there is noopposite pair.Induction case. Let w be a node of height n > 0. As for the base case,all atomic formulas can be forced at w. Non atomic formulas are νβ orπα. Let (w1, w

′1), . . . , (wk, w

′k) be the (non empty) set of successors.

All wi and w′i have a height less than n, so the induction hypothesis

applies to them. Let νβ ∈ (w); due to the construction process of theframe, either νβi

1 ∈ (wi) or νβi2 ∈ (w′

i), for all i = 1, . . . , k, so by theinduction hypothesis, either νβi

1 is forced at wi or νβi2 is forced at w′

i forall i = 1, . . . , k, and so νβ is forced at w. Similarly, let πα ∈ (w); dueto the construction process of the frame, πα1 ∈ (wi) and πα2 ∈ (w′

i)for some i ∈ 1, . . . , k, so by the induction hypothesis, πα1 is forced atwi and πα2 is forced at w′

i for some i, and so πα is forced at w.Lemma II. If Φ is a satisfiable finite set of formulas, then at least one ofthe hypertree-frame of the Φ-auxiliary structure is a (hypertree-)modelof Φ.Proof. We proceed by induction on the complexity of Φ.Base case. If the complexity of Φ is 0, Φ contains only atomic signedformulas, and its hypertree-frame contains the single world w0 whereeach ϕ ∈ Φ is forced; this is possible if and only if Φ is satisfiable. Thisframe clearly is a model of Φ.Induction case. Φ contains at least one non-atomic signed formula.If Φ contains an α-formula, then (Φ\α)∪α1, α2 has lower complexitythan Φ and is also satisfiable; therefore one of the hypertree-frames of(Φ \ α) ∪ α1, α2 is a model of (Φ \ α) ∪ α1, α2, and also of Φ.If Φ contains a β-formula, then (Φ\β)∪β1 or (Φ\β)∪β2 that haslower complexity than Φ, is also satisfiable and has a hypertree-model,which is also a model of Φ.If the set Φ contains only πα-formulas and νβ-formulas,let Φ = P ∪N with P = πα1, . . . , παn and N = νβ1, . . . , νβm.Furthermore, let N12NN = νβ1

iββ1, . . . , νβm

iββm : i1, . . . , im ∈ 1, 2 and νβ ∈

N. If X ∈ N12NN , let X1 be the set of νβ1-elements of X and X2 theset of νβ2-elements of X. If Φ is satisfiable, then for each ordered pair(παi

1, παi2) such that παi ∈ P, there exists an element X ∈ N12NN such

that the sets Φi1 = παi

1 ∪X1 and Φi2 = παi

2 ∪X2 are also satisfiablewith a lower complexity, and therefore have a hypertree-model. Thecorresponding hypertree-model of Φ is obtained as follows : the root isa world w0, with (w0) = Φ, and there are n outgoing hyperarrows,leading to the 2n hypertree-models of the Φi

1,2, i = 1, . . . , n. (Renamingof worlds is used to avoid name clashes.)Comment. The case n = 0 is not ruled out; a set of νβ-formulas isalways B+-satisfiable, and every one-world frame is a model.


Theorem 1. Signed formula ϕ has a model if and only if some hypertree-frame associated with ϕ is a hypertree-model of ϕ.It is an immediate consequence of lemmas I and II.

5.5 Path tree, soundness and completenessLemma III. The path tree associated with a finite (conjunctive) set of

signed formulas is finite.Proof. See §5.3, where the termination of the construction algorithm forthe path tree has been proven.

A line in a hypertree-frame is a sequence (ϕ0, . . . , ϕn) of signed for-mulas such that ϕ0 is the initial formula, ϕi is a child of ϕi−1 and ϕn

has no child.Lemma IV. If S is a path, there exists a hypertree-frame such that Scontains exactly one member of every line.Proof. This is true for the root of the path tree, and if it is true for somepath, it is also true for every successor-path.Lemma V. There is a correspondence between (hypertree-)frames andatomic paths associated with a finite set Φ of formulas; each atomicpath is the set of signed atomic formulas of a frame, and the set ofsigned atomic formulas of each frame is an atomic path.Proof. By induction on the degree of ϕ.

Theorem 2. Signed formula ϕ has a model if and only if (at least) oneof its atomic path does not contain an opposite pair.Proof. Signed formula ϕ has a model if and only if it has a hypertree-model (theorem 1). The corresponding atomic path (lemma V) is con-sistent and does not contain an opposite pair.

Corollary. The method is sound and complete.

Conclusion. We are half-way to an extension of the connection methodto the decidable relevant logic B+. The method introduced here inheritsmost of its properties from the tableau method Bloesch 1993b : sound-ness, completeness and termination. The next step is to obtain a trueconnection method, and to investigate its properties by using a matrix-characterization of B+. Extensions to more powerful systems of relevantlogic should also be possible.

ReferencesBatens D. (2001). “A Dynamic Characterization of the Pure Logic of

Relevant Implication”. Journal of philosophical logic 30 :267–280.Bloesch A. (1993). Signed Tableaux — A Basis for Automated Theorem

Proving in non Classical Logics. PhD thesis, University of Queensland


— (1993). “A Tableau Style Proof System for Two Paraconsistent Log-ics”, Notre Dame Journal of Formal Logic 34:2, pp.295–301.

Fitting M. (1983). Proof Methods for Modal and Intuitionistic Logics.Dordrecht Reidel Publishing Company.

— (1993). “Basic Modal Logic”, Handbook of Logic in Artificial Intelli-gence and Logic Programming ), Ed. by D. M. Gabbay, C. J. Hoggerand J. A. Robinson. Oxford Clarendon Press.

Gore R. (1992).´ Cut-free Sequent and Tableau Systems for PropositionalNormal Logic. Technical report, Cambridge.

Gribomont E.P. & Rossetto D. (1995). CAVEAT: Technique and Toolfor Computer Aided VErification And Transformation, Lecture Notesin Computer Science 939, Computer Aided Verification. Springer.

Hughes G.E. & Cresswell M.J. (1968). An Introduction to Modal Logic.London Methuen and Co Ltd 1968, 1972.

Priest G. & Sylvan R. (1992). “Simplified Semantics for Basic RelevantLogics”, Journal of Philosophical Logic 21:217–232.

Restall G. (1993). “Simplified Semantics for Relevant Logics (and someof their rivals)”, Journal of Philosophical Logic 22:481–511.

Wallen L. (1990). Automated Proof Search in Non-classical Logics. Cam-bridge MIT Press.

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